Unprivileged RISC-V ISA

The RISC-V Instruction Set Manual for CV32A65X: Volume I - Unprivileged Architecture

This document describes the RISC-V unprivileged architecture tailored for OpenHW Group CV32A65X. Not relevant parts (e.g. unsupported extensions) of the original specification are replaced by placeholders.

Contributors to all versions of the spec in alphabetical order (please contact editors to suggest corrections): Derek Atkins, Arvind, Krste Asanović, Rimas Avižienis, Jacob Bachmeyer, Christopher F. Batten, Allen J. Baum, Abel Bernabeu, Alex Bradbury, Scott Beamer, Hans Boehm, Preston Briggs, Christopher Celio, Chuanhua Chang, David Chisnall, Paul Clayton, Palmer Dabbelt, L Peter Deutsch, Ken Dockser, Paul Donahue, Aaron Durbin, Roger Espasa, Greg Favor, Andy Glew, Shaked Flur, Stefan Freudenberger, Marc Gauthier, Andy Glew, Jan Gray, Gianluca Guida, Michael Hamburg, John Hauser, John Ingalls, David Horner, Bruce Hoult, Bill Huffman, Alexandre Joannou, Olof Johansson, Ben Keller, David Kruckemyer, Tariq Kurd, Yunsup Lee, Paul Loewenstein, Daniel Lustig, Yatin Manerkar, Luc Maranget, Ben Marshall, Margaret Martonosi, Phil McCoy, Nathan Menhorn, Christoph Müllner, Joseph Myers, Vijayanand Nagarajan, Rishiyur Nikhil, Jonas Oberhauser, Stefan O’Rear, Markku-Juhani O. Saarinen, Albert Ou, John Ousterhout, Daniel Page, David Patterson, Christopher Pulte, Jose Renau, Josh Scheid, Colin Schmidt, Peter Sewell, Susmit Sarkar, Ved Shanbhogue, Brent Spinney, Brendan Sweeney, Michael Taylor, Wesley Terpstra, Matt Thomas, Tommy Thorn, Philipp Tomsich, Caroline Trippel, Ray VanDeWalker, Muralidaran Vijayaraghavan, Megan Wachs, Paul Wamsley, Andrew Waterman, Robert Watson, David Weaver, Derek Williams, Claire Wolf, Andrew Wright, Reinoud Zandijk, and Sizhuo Zhang.

This document is released under a Creative Commons Attribution 4.0 International License.

This document is a derivative of “The RISC-V Instruction Set Manual, Volume I: User-Level ISA Version 2.1” released under the following license: ©2010-2017 Andrew Waterman, Yunsup Lee, David Patterson, Krste Asanović. Creative Commons Attribution 4.0 International License. Please cite as: “The RISC-V Instruction Set Manual, Volume I: User-Level ISA, Document Version 20191214-draft”, Editors Andrew Waterman and Krste Asanović, RISC-V Foundation, December 2019.

Contributors to CV32A65X versions of the spec in alphabetical order: Jean-Roch Coulon, André Sintzoff.

Preface

Preface to Document Version for CV32A65X

This document describes the RISC-V unprivileged architecture tailored for OpenHW Group CV32A65X.

Preface to Document Version 20241017

This document describes the RISC-V unprivileged architecture.

The ISA modules marked Ratified have been ratified at this time. The modules marked Frozen are not expected to change significantly before being put up for ratification. The modules marked Draft are expected to change before ratification.

The document contains the following versions of the RISC-V ISA modules:

Base Version Status

RV32I

2.1

Ratified

RV32E

2.0

Ratified

RV64E

2.0

Ratified

RV64I

2.1

Ratified

RV128I

1.7

Draft

Extension

Version

Status

Zifencei

2.0

Ratified

Zicsr

2.0

Ratified

Zicntr

2.0

Ratified

Zihintntl

1.0

Ratified

Zihintpause

2.0

Ratified

Zimop

1.0

Ratified

Zicond

1.0

Ratified

M

2.0

Ratified

Zmmul

1.0

Ratified

A

2.1

Ratified

Zawrs

1.01

Ratified

Zacas

1.0

Ratified

Zabha

1.0

Ratified

RVWMO

2.0

Ratified

Ztso

1.0

Ratified

CMO

1.0

Ratified

F

2.2

Ratified

D

2.2

Ratified

Q

2.2

Ratified

Zfh

1.0

Ratified

Zfhmin

1.0

Ratified

Zfa

1.0

Ratified

Zfinx

1.0

Ratified

Zdinx

1.0

Ratified

Zhinx

1.0

Ratified

Zhinxmin

1.0

Ratified

C

2.0

Ratified

Zce

1.0

Ratified

B

1.0

Ratified

P

0.2

Draft

V

1.0

Ratified

Zbkb

1.0

Ratified

Zbkc

1.0

Ratified

Zbkx

1.0

Ratified

Zk

1.0

Ratified

Zks

1.0

Ratified

Zvbb

1.0

Ratified

Zvbc

1.0

Ratified

Zvkg

1.0

Ratified

Zvkned

1.0

Ratified

Zvknhb

1.0

Ratified

Zvksed

1.0

Ratified

Zvksh

1.0

Ratified

Zvkt

1.0

Ratified

Zicfiss

1.0

Ratified

Zicfilp

1.0

Ratified

The changes in this version of the document include:

  • The inclusion of all ratified extensions through March 2024.

  • The draft Zam extension has been removed, in favor of the definition of a misaligned atomicity granule PMA.

  • The concept of vacant memory regions has been superseded by inaccessible memory or I/O regions.

Preface to Document Version 20191213-Base-Ratified

This document describes the RISC-V unprivileged architecture.

The ISA modules marked Ratified have been ratified at this time. The modules marked Frozen are not expected to change significantly before being put up for ratification. The modules marked Draft are expected to change before ratification.

The document contains the following versions of the RISC-V ISA modules:

Base Version Status

RVWMO

2.0

Ratified

RV32I

2.1

Ratified

RV64I

2.1

Ratified

RV32E

1.9

Draft

RV128I

1.7

Draft

Extension

Version

Status

M

2.0

Ratified

A

2.1

Ratified

F

2.2

Ratified

D

2.2

Ratified

Q

2.2

Ratified

C

2.0

Ratified

Counters

2.0

Draft

L

0.0

Draft

B

0.0

Draft

J

0.0

Draft

T

0.0

Draft

P

0.2

Draft

V

0.7

Draft

Zicsr

2.0

Ratified

Zifencei

2.0

Ratified

Zam

0.1

Draft

Ztso

0.1

Frozen

The changes in this version of the document include:

  • The A extension, now version 2.1, was ratified by the board in December 2019.

  • Defined big-endian ISA variant.

  • Moved N extension for user-mode interrupts into Volume II.

  • Defined PAUSE hint instruction.

Preface to Document Version 20190608-Base-Ratified

This document describes the RISC-V unprivileged architecture.

The RVWMO memory model has been ratified at this time. The ISA modules marked Ratified, have been ratified at this time. The modules marked Frozen are not expected to change significantly before being put up for ratification. The modules marked Draft are expected to change before ratification.

The document contains the following versions of the RISC-V ISA modules:

Base Version Status

RVWMO

2.0

Ratified

RV32I

2.1

Ratified

RV64I

2.1

Ratified

RV32E

1.9

Draft

RV128I

1.7

Draft

Extension

Version

Status

Zifencei

2.0

Ratified

Zicsr

2.0

Ratified

M

2.0

Ratified

A

2.0

Frozen

F

2.2

Ratified

D

2.2

Ratified

Q

2.2

Ratified

C

2.0

Ratified

Ztso

0.1

Frozen

Counters

2.0

Draft

L

0.0

Draft

B

0.0

Draft

J

0.0

Draft

T

0.0

Draft

P

0.2

Draft

V

0.7

Draft

Zam

0.1

Draft

The changes in this version of the document include:

  • Moved description to Ratified for the ISA modules ratified by the board in early 2019.

  • Removed the A extension from ratification.

  • Changed document version scheme to avoid confusion with versions of the ISA modules.

  • Incremented the version numbers of the base integer ISA to 2.1, reflecting the presence of the ratified RVWMO memory model and exclusion of FENCE.I, counters, and CSR instructions that were in previous base ISA.

  • Incremented the version numbers of the F and D extensions to 2.2, reflecting that version 2.1 changed the canonical NaN, and version 2.2 defined the NaN-boxing scheme and changed the definition of the FMIN and FMAX instructions.

  • Changed name of document to refer to "unprivileged" instructions as part of move to separate ISA specifications from platform profile mandates.

  • Added clearer and more precise definitions of execution environments, harts, traps, and memory accesses.

  • Defined instruction-set categories: standard, reserved, custom, non-standard, and non-conforming.

  • Removed text implying operation under alternate endianness, as alternate-endianness operation has not yet been defined for RISC-V.

  • Changed description of misaligned load and store behavior. The specification now allows visible misaligned address traps in execution environment interfaces, rather than just mandating invisible handling of misaligned loads and stores in user mode. Also, now allows access-fault exceptions to be reported for misaligned accesses (including atomics) that should not be emulated.

  • Moved FENCE.I out of the mandatory base and into a separate extension, with Zifencei ISA name. FENCE.I was removed from the Linux user ABI and is problematic in implementations with large incoherent instruction and data caches. However, it remains the only standard instruction-fetch coherence mechanism.

  • Removed prohibitions on using RV32E with other extensions.

  • Removed platform-specific mandates that certain encodings produce illegal-instruction exceptions in RV32E and RV64I chapters.

  • Counter/timer instructions are now not considered part of the mandatory base ISA, and so CSR instructions were moved into separate chapter and marked as version 2.0, with the unprivileged counters moved into another separate chapter. The counters are not ready for ratification as there are outstanding issues, including counter inaccuracies.

  • A CSR-access ordering model has been added.

  • Explicitly defined the 16-bit half-precision floating-point format for floating-point instructions in the 2-bit fmt field.

  • Defined the signed-zero behavior of FMIN.fmt and FMAX.fmt, and changed their behavior on signaling-NaN inputs to conform to the minimumNumber and maximumNumber operations in the proposed IEEE 754-201x specification.

  • The memory consistency model, RVWMO, has been defined.

  • The "Zam" extension, which permits misaligned AMOs and specifies their semantics, has been defined.

  • The "Ztso" extension, which enforces a stricter memory consistency model than RVWMO, has been defined.

  • Improvements to the description and commentary.

  • Defined the term IALIGN as shorthand to describe the instruction-address alignment constraint.

  • Removed text of P extension chapter as now superseded by active task group documents.

  • Removed text of V extension chapter as now superseded by separate vector extension draft document.

Preface to Document Version 2.2

This is version 2.2 of the document describing the RISC-V user-level architecture. The document contains the following versions of the RISC-V ISA modules:

Base Version Draft Frozen?

RV32I

2.0

Y

RV32E

1.9

N

RV64I

2.0

Y

RV128I

1.7

N

Extension

Version

Frozen?

M

2.0

Y

A

2.0

Y

F

2.0

Y

D

2.0

Y

Q

2.0

Y

L

0.0

N

C

2.0

Y

B

0.0

N

J

0.0

N

T

0.0

N

P

0.1

N

V

0.7

N

N

1.1

N

To date, no parts of the standard have been officially ratified by the RISC-V Foundation, but the components labeled "frozen" above are not expected to change during the ratification process beyond resolving ambiguities and holes in the specification.

The major changes in this version of the document include:

  • The previous version of this document was released under a Creative Commons Attribution 4.0 International License by the original authors, and this and future versions of this document will be released under the same license.

  • Rearranged chapters to put all extensions first in canonical order.

  • Improvements to the description and commentary.

  • Modified implicit hinting suggestion on JALR to support more efficient macro-op fusion of LUI/JALR and AUIPC/JALR pairs.

  • Clarification of constraints on load-reserved/store-conditional sequences.

  • A new table of control and status register (CSR) mappings.

  • Clarified purpose and behavior of high-order bits of fcsr.

  • Corrected the description of the FNMADD.fmt and FNMSUB.fmt instructions, which had suggested the incorrect sign of a zero result.

  • Instructions FMV.S.X and FMV.X.S were renamed to FMV.W.X and FMV.X.W respectively to be more consistent with their semantics, which did not change. The old names will continue to be supported in the tools.

  • Specified behavior of narrower (stem cec1626943fd22b8bc5d3763224c7c1fFLEN) floating-point values held in wider f registers using NaN-boxing model.

  • Defined the exception behavior of FMA(stem 66ad238aa9cc3ef0bde4119bf3f5e449, 0, qNaN).

  • Added note indicating that the P extension might be reworked into an integer packed-SIMD proposal for fixed-point operations using the integer registers.

  • A draft proposal of the V vector instruction-set extension.

  • An early draft proposal of the N user-level traps extension.

  • An expanded pseudoinstruction listing.

  • Removal of the calling convention chapter, which has been superseded by the RISC-V ELF psABI Specification (RISC-V ELF PsABI Specification, n.d.).

  • The C extension has been frozen and renumbered version 2.0.

Preface to Document Version 2.1

This is version 2.1 of the document describing the RISC-V user-level architecture. Note the frozen user-level ISA base and extensions IMAFDQ version 2.0 have not changed from the previous version of this document (Waterman et al., 2014), but some specification holes have been fixed and the documentation has been improved. Some changes have been made to the software conventions.

  • Numerous additions and improvements to the commentary sections.

  • Separate version numbers for each chapter.

  • Modification to long instruction encodings stem d31e2e21d71d81f6406a849380df264164 bits to avoid moving the rd specifier in very long instruction formats.

  • CSR instructions are now described in the base integer format where the counter registers are introduced, as opposed to only being introduced later in the floating-point section (and the companion privileged architecture manual).

  • The SCALL and SBREAK instructions have been renamed to ECALL and EBREAK, respectively. Their encoding and functionality are unchanged.

  • Clarification of floating-point NaN handling, and a new canonical NaN value.

  • Clarification of values returned by floating-point to integer conversions that overflow.

  • Clarification of LR/SC allowed successes and required failures, including use of compressed instructions in the sequence.

  • A new RV32E base ISA proposal for reduced integer register counts, supports MAC extensions.

  • A revised calling convention.

  • Relaxed stack alignment for soft-float calling convention, and description of the RV32E calling convention.

  • A revised proposal for the C compressed extension, version 1.9 .

Preface to Version 2.0

This is the second release of the user ISA specification, and we intend the specification of the base user ISA plus general extensions (i.e., IMAFD) to remain fixed for future development. The following changes have been made since Version 1.0 (Waterman et al., 2011) of this ISA specification.

  • The ISA has been divided into an integer base with several standard extensions.

  • The instruction formats have been rearranged to make immediate encoding more efficient.

  • The base ISA has been defined to have a little-endian memory system, with big-endian or bi-endian as non-standard variants.

  • Load-Reserved/Store-Conditional (LR/SC) instructions have been added in the atomic instruction extension.

  • AMOs and LR/SC can support the release consistency model.

  • The FENCE instruction provides finer-grain memory and I/O orderings.

  • An AMO for fetch-and-XOR (AMOXOR) has been added, and the encoding for AMOSWAP has been changed to make room.

  • The AUIPC instruction, which adds a 20-bit upper immediate to the PC, replaces the RDNPC instruction, which only read the current PC value. This results in significant savings for position-independent code.

  • The JAL instruction has now moved to the U-Type format with an explicit destination register, and the J instruction has been dropped being replaced by JAL with rd=x0. This removes the only instruction with an implicit destination register and removes the J-Type instruction format from the base ISA. There is an accompanying reduction in JAL reach, but a significant reduction in base ISA complexity.

  • The static hints on the JALR instruction have been dropped. The hints are redundant with the rd and rs1 register specifiers for code compliant with the standard calling convention.

  • The JALR instruction now clears the lowest bit of the calculated target address, to simplify hardware and to allow auxiliary information to be stored in function pointers.

  • The MFTX.S and MFTX.D instructions have been renamed to FMV.X.S and FMV.X.D, respectively. Similarly, MXTF.S and MXTF.D instructions have been renamed to FMV.S.X and FMV.D.X, respectively.

  • The MFFSR and MTFSR instructions have been renamed to FRCSR and FSCSR, respectively. FRRM, FSRM, FRFLAGS, and FSFLAGS instructions have been added to individually access the rounding mode and exception flags subfields of the fcsr.

  • The FMV.X.S and FMV.X.D instructions now source their operands from rs1, instead of rs2. This change simplifies datapath design.

  • FCLASS.S and FCLASS.D floating-point classify instructions have been added.

  • A simpler NaN generation and propagation scheme has been adopted.

  • For RV32I, the system performance counters have been extended to 64-bits wide, with separate read access to the upper and lower 32 bits.

  • Canonical NOP and MV encodings have been defined.

  • Standard instruction-length encodings have been defined for 48-bit, 64-bit, and stem d31e2e21d71d81f6406a849380df264164-bit instructions.

  • Description of a 128-bit address space variant, RV128, has been added.

  • Major opcodes in the 32-bit base instruction format have been allocated for user-defined custom extensions.

  • A typographical error that suggested that stores source their data from rd has been corrected to refer to rs2.

1. Introduction

RISC-V (pronounced "risk-five") is a new instruction-set architecture (ISA) that was originally designed to support computer architecture research and education, but which we now hope will also become a standard free and open architecture for industry implementations. Our goals in defining RISC-V include:

  • A completely open ISA that is freely available to academia and industry.

  • A real ISA suitable for direct native hardware implementation, not just simulation or binary translation.

  • An ISA that avoids "over-architecting" for a particular microarchitecture style (e.g., microcoded, in-order, decoupled, out-of-order) or implementation technology (e.g., full-custom, ASIC, FPGA), but which allows efficient implementation in any of these.

  • An ISA separated into a small base integer ISA, usable by itself as a base for customized accelerators or for educational purposes, and optional standard extensions, to support general-purpose software development.

  • Support for the revised 2008 IEEE-754 floating-point standard. (ANSI/IEEE Std 754-2008, IEEE Standard for Floating-Point Arithmetic, 2008)

  • An ISA supporting extensive ISA extensions and specialized variants.

  • Both 32-bit and 64-bit address space variants for applications, operating system kernels, and hardware implementations.

  • An ISA with support for highly parallel multicore or manycore implementations, including heterogeneous multiprocessors.

  • Optional variable-length instructions to both expand available instruction encoding space and to support an optional dense instruction encoding for improved performance, static code size, and energy efficiency.

  • A fully virtualizable ISA to ease hypervisor development.

  • An ISA that simplifies experiments with new privileged architecture designs.

Commentary on our design decisions is formatted as in this paragraph. This non-normative text can be skipped if the reader is only interested in the specification itself.

The name RISC-V was chosen to represent the fifth major RISC ISA design from UC Berkeley (RISC-I (Patterson & Séquin, 1981), RISC-II (Katevenis et al., 1983), SOAR (Ungar et al., 1984), and SPUR (Lee et al., 1989) were the first four). We also pun on the use of the Roman numeral "V" to signify "variations" and "vectors", as support for a range of architecture research, including various data-parallel accelerators, is an explicit goal of the ISA design.

The RISC-V ISA is defined avoiding implementation details as much as possible (although commentary is included on implementation-driven decisions) and should be read as the software-visible interface to a wide variety of implementations rather than as the design of a particular hardware artifact. The RISC-V manual is structured in two volumes. This volume covers the design of the base unprivileged instructions, including optional unprivileged ISA extensions. Unprivileged instructions are those that are generally usable in all privilege modes in all privileged architectures, though behavior might vary depending on privilege mode and privilege architecture. The second volume provides the design of the first ("classic") privileged architecture. The manuals use IEC 80000-13:2008 conventions, with a byte of 8 bits.

In the unprivileged ISA design, we tried to remove any dependence on particular microarchitectural features, such as cache line size, or on privileged architecture details, such as page translation. This is both for simplicity and to allow maximum flexibility for alternative microarchitectures or alternative privileged architectures.

1.1. RISC-V Hardware Platform Terminology

A RISC-V hardware platform can contain one or more RISC-V-compatible processing cores together with other non-RISC-V-compatible cores, fixed-function accelerators, various physical memory structures, I/O devices, and an interconnect structure to allow the components to communicate.

A component is termed a core if it contains an independent instruction fetch unit. A RISC-V-compatible core might support multiple RISC-V-compatible hardware threads, or harts, through multithreading.

A RISC-V core might have additional specialized instruction-set extensions or an added coprocessor. We use the term coprocessor to refer to a unit that is attached to a RISC-V core and is mostly sequenced by a RISC-V instruction stream, but which contains additional architectural state and instruction-set extensions, and possibly some limited autonomy relative to the primary RISC-V instruction stream.

We use the term accelerator to refer to either a non-programmable fixed-function unit or a core that can operate autonomously but is specialized for certain tasks. In RISC-V systems, we expect many programmable accelerators will be RISC-V-based cores with specialized instruction-set extensions and/or customized coprocessors. An important class of RISC-V accelerators are I/O accelerators, which offload I/O processing tasks from the main application cores.

The system-level organization of a RISC-V hardware platform can range from a single-core microcontroller to a many-thousand-node cluster of shared-memory manycore server nodes. Even small systems-on-a-chip might be structured as a hierarchy of multicomputers and/or multiprocessors to modularize development effort or to provide secure isolation between subsystems.

1.2. RISC-V Software Execution Environments and Harts

The behavior of a RISC-V program depends on the execution environment in which it runs. A RISC-V execution environment interface (EEI) defines the initial state of the program, the number and type of harts in the environment including the privilege modes supported by the harts, the accessibility and attributes of memory and I/O regions, the behavior of all legal instructions executed on each hart (i.e., the ISA is one component of the EEI), and the handling of any interrupts or exceptions raised during execution including environment calls. Examples of EEIs include the Linux application binary interface (ABI), or the RISC-V supervisor binary interface (SBI). The implementation of a RISC-V execution environment can be pure hardware, pure software, or a combination of hardware and software. For example, opcode traps and software emulation can be used to implement functionality not provided in hardware. Examples of execution environment implementations include:

  • "Bare metal" hardware platforms where harts are directly implemented by physical processor threads and instructions have full access to the physical address space. The hardware platform defines an execution environment that begins at power-on reset.

  • RISC-V operating systems that provide multiple user-level execution environments by multiplexing user-level harts onto available physical processor threads and by controlling access to memory via virtual memory.

  • RISC-V hypervisors that provide multiple supervisor-level execution environments for guest operating systems.

  • RISC-V emulators, such as Spike, QEMU or rv8, which emulate RISC-V harts on an underlying x86 system, and which can provide either a user-level or a supervisor-level execution environment.

A bare hardware platform can be considered to define an EEI, where the accessible harts, memory, and other devices populate the environment, and the initial state is that at power-on reset. Generally, most software is designed to use a more abstract interface to the hardware, as more abstract EEIs provide greater portability across different hardware platforms. Often EEIs are layered on top of one another, where one higher-level EEI uses another lower-level EEI.

From the perspective of software running in a given execution environment, a hart is a resource that autonomously fetches and executes RISC-V instructions within that execution environment. In this respect, a hart behaves like a hardware thread resource even if time-multiplexed onto real hardware by the execution environment. Some EEIs support the creation and destruction of additional harts, for example, via environment calls to fork new harts.

The execution environment is responsible for ensuring the eventual forward progress of each of its harts. For a given hart, that responsibility is suspended while the hart is exercising a mechanism that explicitly waits for an event, such as the wait-for-interrupt instruction defined in Volume II of this specification; and that responsibility ends if the hart is terminated. The following events constitute forward progress:

  • The retirement of an instruction.

  • A trap, as defined in Section 1.6.

  • Any other event defined by an extension to constitute forward progress.

The term hart was introduced in the work on Lithe (Pan et al., 2009) and (Pan et al., 2010) to provide a term to represent an abstract execution resource as opposed to a software thread programming abstraction.

The important distinction between a hardware thread (hart) and a software thread context is that the software running inside an execution environment is not responsible for causing progress of each of its harts; that is the responsibility of the outer execution environment. So the environment’s harts operate like hardware threads from the perspective of the software inside the execution environment.

An execution environment implementation might time-multiplex a set of guest harts onto fewer host harts provided by its own execution environment but must do so in a way that guest harts operate like independent hardware threads. In particular, if there are more guest harts than host harts then the execution environment must be able to preempt the guest harts and must not wait indefinitely for guest software on a guest hart to "yield" control of the guest hart.

1.3. RISC-V ISA Overview

A RISC-V ISA is defined as a base integer ISA, which must be present in any implementation, plus optional extensions to the base ISA. The base integer ISAs are very similar to that of the early RISC processors except with no branch delay slots and with support for optional variable-length instruction encodings. A base is carefully restricted to a minimal set of instructions sufficient to provide a reasonable target for compilers, assemblers, linkers, and operating systems (with additional privileged operations), and so provides a convenient ISA and software toolchain "skeleton" around which more customized processor ISAs can be built.

Although it is convenient to speak of the RISC-V ISA, RISC-V is actually a family of related ISAs, of which there are currently four base ISAs. Each base integer instruction set is characterized by the width of the integer registers and the corresponding size of the address space and by the number of integer registers. There are two primary base integer variants, RV32I and RV64I, described in Chapter 2 and Chapter 4, which provide 32-bit or 64-bit address spaces respectively. We use the term XLEN to refer to the width of an integer register in bits (either 32 or 64). Chapter 3 describes the RV32E and RV64E subset variants of the RV32I or RV64I base instruction sets respectively, which have been added to support small microcontrollers, and which have half the number of integer registers. Chapter 5 sketches a future RV128I variant of the base integer instruction set supporting a flat 128-bit address space (XLEN=128). The base integer instruction sets use a two’s-complement representation for signed integer values.

Although 64-bit address spaces are a requirement for larger systems, we believe 32-bit address spaces will remain adequate for many embedded and client devices for decades to come and will be desirable to lower memory traffic and energy consumption. In addition, 32-bit address spaces are sufficient for educational purposes. A larger flat 128-bit address space might eventually be required, so we ensured this could be accommodated within the RISC-V ISA framework.

The four base ISAs in RISC-V are treated as distinct base ISAs. A common question is why is there not a single ISA, and in particular, why is RV32I not a strict subset of RV64I? Some earlier ISA designs (SPARC, MIPS) adopted a strict superset policy when increasing address space size to support running existing 32-bit binaries on new 64-bit hardware.

The main advantage of explicitly separating base ISAs is that each base ISA can be optimized for its needs without requiring to support all the operations needed for other base ISAs. For example, RV64I can omit instructions and CSRs that are only needed to cope with the narrower registers in RV32I. The RV32I variants can use encoding space otherwise reserved for instructions only required by wider address-space variants.

The main disadvantage of not treating the design as a single ISA is that it complicates the hardware needed to emulate one base ISA on another (e.g., RV32I on RV64I). However, differences in addressing and illegal-instruction traps generally mean some mode switch would be required in hardware in any case even with full superset instruction encodings, and the different RISC-V base ISAs are similar enough that supporting multiple versions is relatively low cost. Although some have proposed that the strict superset design would allow legacy 32-bit libraries to be linked with 64-bit code, this is impractical in practice, even with compatible encodings, due to the differences in software calling conventions and system-call interfaces.

The RISC-V privileged architecture provides fields in misa to control the unprivileged ISA at each level to support emulating different base ISAs on the same hardware. We note that newer SPARC and MIPS ISA revisions have deprecated support for running 32-bit code unchanged on 64-bit systems.

A related question is why there is a different encoding for 32-bit adds in RV32I (ADD) and RV64I (ADDW)? The ADDW opcode could be used for 32-bit adds in RV32I and ADDD for 64-bit adds in RV64I, instead of the existing design which uses the same opcode ADD for 32-bit adds in RV32I and 64-bit adds in RV64I with a different opcode ADDW for 32-bit adds in RV64I. This would also be more consistent with the use of the same LW opcode for 32-bit load in both RV32I and RV64I. The very first versions of RISC-V ISA did have a variant of this alternate design, but the RISC-V design was changed to the current choice in January 2011. Our focus was on supporting 32-bit integers in the 64-bit ISA not on providing compatibility with the 32-bit ISA, and the motivation was to remove the asymmetry that arose from having not all opcodes in RV32I have a *W suffix (e.g., ADDW, but AND not ANDW). In hindsight, this was perhaps not well-justified and a consequence of designing both ISAs at the same time as opposed to adding one later to sit on top of another, and also from a belief we had to fold platform requirements into the ISA spec which would imply that all the RV32I instructions would have been required in RV64I. It is too late to change the encoding now, but this is also of little practical consequence for the reasons stated above.

It has been noted we could enable the *W variants as an extension to RV32I systems to provide a common encoding across RV64I and a future RV32 variant.

RISC-V has been designed to support extensive customization and specialization. Each base integer ISA can be extended with one or more optional instruction-set extensions. An extension may be categorized as either standard, custom, or non-conforming. For this purpose, we divide each RISC-V instruction-set encoding space (and related encoding spaces such as the CSRs) into three disjoint categories: standard, reserved, and custom. Standard extensions and encodings are defined by RISC-V International; any extensions not defined by RISC-V International are non-standard. Each base ISA and its standard extensions use only standard encodings, and shall not conflict with each other in their uses of these encodings. Reserved encodings are currently not defined but are saved for future standard extensions; once thus used, they become standard encodings. Custom encodings shall never be used for standard extensions and are made available for vendor-specific non-standard extensions. Non-standard extensions are either custom extensions, that use only custom encodings, or non-conforming extensions, that use any standard or reserved encoding. Instruction-set extensions are generally shared but may provide slightly different functionality depending on the base ISA. Chapter 37 describes various ways of extending the RISC-V ISA. We have also developed a naming convention for RISC-V base instructions and instruction-set extensions, described in detail in Chapter 38.

To support more general software development, a set of standard extensions are defined to provide integer multiply/divide, atomic operations, and single and double-precision floating-point arithmetic. The base integer ISA is named "I" (prefixed by RV32 or RV64 depending on integer register width), and contains integer computational instructions, integer loads, integer stores, and control-flow instructions. The standard integer multiplication and division extension is named "M", and adds instructions to multiply and divide values held in the integer registers. The standard atomic instruction extension, denoted by "A", adds instructions that atomically read, modify, and write memory for inter-processor synchronization. The standard single-precision floating-point extension, denoted by "F", adds floating-point registers, single-precision computational instructions, and single-precision loads and stores. The standard double-precision floating-point extension, denoted by "D", expands the floating-point registers, and adds double-precision computational instructions, loads, and stores. The standard "C" compressed instruction extension provides narrower 16-bit forms of common instructions.

Beyond the base integer ISA and these standard extensions, we believe it is rare that a new instruction will provide a significant benefit for all applications, although it may be very beneficial for a certain domain. As energy efficiency concerns are forcing greater specialization, we believe it is important to simplify the required portion of an ISA specification. Whereas other architectures usually treat their ISA as a single entity, which changes to a new version as instructions are added over time, RISC-V will endeavor to keep the base and each standard extension constant over time, and instead layer new instructions as further optional extensions. For example, the base integer ISAs will continue as fully supported standalone ISAs, regardless of any subsequent extensions.

1.4. Memory

A RISC-V hart has a single byte-addressable address space of stem d27f663ab9e7364c59a70b83260dca56 bytes for all memory accesses. A word of memory is defined as 32 bits (4 bytes). Correspondingly, a halfword is 16 bits (2 bytes), a doubleword is 64 bits (8 bytes), and a quadword is 128 bits (16 bytes). The memory address space is circular, so that the byte at address stem 52d22d5bedfb57a75620c1335aa01e21 is adjacent to the byte at address zero. Accordingly, memory address computations done by the hardware ignore overflow and instead wrap around modulo stem d27f663ab9e7364c59a70b83260dca56.

The execution environment determines the mapping of hardware resources into a hart’s address space. Different address ranges of a hart’s address space may (1) contain main memory, or (2) contain one or more I/O devices. Reads and writes of I/O devices may have visible side effects, but accesses to main memory cannot. Vacant address ranges are not a separate category but can be represented as either main memory or I/O regions that are not accessible. Although it is possible for the execution environment to call everything in a hart’s address space an I/O device, it is usually expected that some portion will be specified as main memory.

When a RISC-V platform has multiple harts, the address spaces of any two harts may be entirely the same, or entirely different, or may be partly different but sharing some subset of resources, mapped into the same or different address ranges.

For a purely "bare metal" environment, all harts may see an identical address space, accessed entirely by physical addresses. However, when the execution environment includes an operating system employing address translation, it is common for each hart to be given a virtual address space that is largely or entirely its own.

Executing each RISC-V machine instruction entails one or more memory accesses, subdivided into implicit and explicit accesses. For each instruction executed, an implicit memory read (instruction fetch) is done to obtain the encoded instruction to execute. Many RISC-V instructions perform no further memory accesses beyond instruction fetch. Specific load and store instructions perform an explicit read or write of memory at an address determined by the instruction. The execution environment may dictate that instruction execution performs other implicit memory accesses (such as to implement address translation) beyond those documented for the unprivileged ISA.

The execution environment determines what portions of the address space are accessible for each kind of memory access. For example, the set of locations that can be implicitly read for instruction fetch may or may not have any overlap with the set of locations that can be explicitly read by a load instruction; and the set of locations that can be explicitly written by a store instruction may be only a subset of locations that can be read. Ordinarily, if an instruction attempts to access memory at an inaccessible address, an exception is raised for the instruction.

Except when specified otherwise, implicit reads that do not raise an exception and that have no side effects may occur arbitrarily early and speculatively, even before the machine could possibly prove that the read will be needed. For instance, a valid implementation could attempt to read all of main memory at the earliest opportunity, cache as many fetchable (executable) bytes as possible for later instruction fetches, and avoid reading main memory for instruction fetches ever again. To ensure that certain implicit reads are ordered only after writes to the same memory locations, software must execute specific fence or cache-control instructions defined for this purpose (such as the FENCE.I instruction defined in Chapter 6).

The memory accesses (implicit or explicit) made by a hart may appear to occur in a different order as perceived by another hart or by any other agent that can access the same memory. This perceived reordering of memory accesses is always constrained, however, by the applicable memory consistency model. The default memory consistency model for RISC-V is the RISC-V Weak Memory Ordering (RVWMO), defined in Chapter 18 and in appendices. Optionally, an implementation may adopt the stronger model of Total Store Ordering, as defined in Chapter 19. The execution environment may also add constraints that further limit the perceived reordering of memory accesses. Since the RVWMO model is the weakest model allowed for any RISC-V implementation, software written for this model is compatible with the actual memory consistency rules of all RISC-V implementations. As with implicit reads, software must execute fence or cache-control instructions to ensure specific ordering of memory accesses beyond the requirements of the assumed memory consistency model and execution environment.

1.5. Base Instruction-Length Encoding

The base RISC-V ISA has fixed-length 32-bit instructions that must be naturally aligned on 32-bit boundaries. However, the standard RISC-V encoding scheme is designed to support ISA extensions with variable-length instructions, where each instruction can be any number of 16-bit instruction parcels in length and parcels are naturally aligned on 16-bit boundaries. The standard compressed ISA extension described in Chapter 28 reduces code size by providing compressed 16-bit instructions and relaxes the alignment constraints to allow all instructions (16 bit and 32 bit) to be aligned on any 16-bit boundary to improve code density.

We use the term IALIGN (measured in bits) to refer to the instruction-address alignment constraint the implementation enforces. IALIGN is 32 bits in the base ISA, but some ISA extensions, including the compressed ISA extension, relax IALIGN to 16 bits. IALIGN may not take on any value other than 16 or 32.

We use the term ILEN (measured in bits) to refer to the maximum instruction length supported by an implementation, and which is always a multiple of IALIGN. For implementations supporting only a base instruction set, ILEN is 32 bits. Implementations supporting longer instructions have larger values of ILEN.

Table 1 illustrates the standard RISC-V instruction-length encoding convention. All the 32-bit instructions in the base ISA have their lowest two bits set to 11. The optional compressed 16-bit instruction-set extensions have their lowest two bits equal to 00, 01, or 10.

1.5.1. Expanded Instruction-Length Encoding

A portion of the 32-bit instruction-encoding space has been tentatively allocated for instructions longer than 32 bits. The entirety of this space is reserved at this time, and the following proposal for encoding instructions longer than 32 bits is not considered frozen.

Standard instruction-set extensions encoded with more than 32 bits have additional low-order bits set to 1, with the conventions for 48-bit and 64-bit lengths shown in Table 1. Instruction lengths between 80 bits and 176 bits are encoded using a 3-bit field in bits [14:12] giving the number of 16-bit words in addition to the first 5stem 715efe442a3a591fd63a05098adb4ab716-bit words. The encoding with bits [14:12] set to "111" is reserved for future longer instruction encodings.

Table 1. RISC-V instruction length encoding. Only the 16-bit and 32-bit encodings are considered frozen at this time.
xxxxxxxxxxxxxxaa 16-bit (aa≠11)

xxxxxxxxxxxxxxxx

xxxxxxxxxxxbbb11

32-bit (bbb≠111)

stem 56ee3cdc7603fdb6d68d8551a34e9a36xxxx

xxxxxxxxxxxxxxxx

xxxxxxxxxx011111

48-bit

stem 56ee3cdc7603fdb6d68d8551a34e9a36xxxx

xxxxxxxxxxxxxxxx

xxxxxxxxx0111111

64-bit

stem 56ee3cdc7603fdb6d68d8551a34e9a36xxxx

xxxxxxxxxxxxxxxx

xnnnxxxxx1111111

(80+16*nnn)-bit, nnn≠111

stem 56ee3cdc7603fdb6d68d8551a34e9a36xxxx

xxxxxxxxxxxxxxxx

x111xxxxx1111111

Reserved for ≥192-bits

Byte Address:

base+4

base+2

base

Given the code size and energy savings of a compressed format, we wanted to build in support for a compressed format to the ISA encoding scheme rather than adding this as an afterthought, but to allow simpler implementations we didn’t want to make the compressed format mandatory. We also wanted to optionally allow longer instructions to support experimentation and larger instruction-set extensions. Although our encoding convention required a tighter encoding of the core RISC-V ISA, this has several beneficial effects.

An implementation of the standard IMAFD ISA need only hold the most-significant 30 bits in instruction caches (a 6.25% saving). On instruction cache refills, any instructions encountered with either low bit clear should be recoded into illegal 30-bit instructions before storing in the cache to preserve illegal-instruction exception behavior.

Perhaps more importantly, by condensing our base ISA into a subset of the 32-bit instruction word, we leave more space available for non-standard and custom extensions. In particular, the base RV32I ISA uses less than 1/8 of the encoding space in the 32-bit instruction word. As described in Chapter 37, an implementation that does not require support for the standard compressed instruction extension can map 3 additional non-conforming 30-bit instruction spaces into the 32-bit fixed-width format, while preserving support for standard ≥32-bit instruction-set extensions. Further, if the implementation also does not need instructions >32-bits in length, it can recover a further four major opcodes for non-conforming extensions.

Encodings with bits [15:0] all zeros are defined as illegal instructions. These instructions are considered to be of minimal length: 16 bits if any 16-bit instruction-set extension is present, otherwise 32 bits. The encoding with bits [ILEN-1:0] all ones is also illegal; this instruction is considered to be ILEN bits long.

We consider it a feature that any length of instruction containing all zero bits is not legal, as this quickly traps erroneous jumps into zeroed memory regions. Similarly, we also reserve the instruction encoding containing all ones to be an illegal instruction, to catch the other common pattern observed with unprogrammed non-volatile memory devices, disconnected memory buses, or broken memory devices.

Software can rely on a naturally aligned 32-bit word containing zero to act as an illegal instruction on all RISC-V implementations, to be used by software where an illegal instruction is explicitly desired. Defining a corresponding known illegal value for all ones is more difficult due to the variable-length encoding. Software cannot generally use the illegal value of ILEN bits of all 1s, as software might not know ILEN for the eventual target machine (e.g., if software is compiled into a standard binary library used by many different machines). Defining a 32-bit word of all ones as illegal was also considered, as all machines must support a 32-bit instruction size, but this requires the instruction-fetch unit on machines with ILEN >32 report an illegal-instruction exception rather than an access-fault exception when such an instruction borders a protection boundary, complicating variable-instruction-length fetch and decode.

RISC-V base ISAs have either little-endian or big-endian memory systems, with the privileged architecture further defining bi-endian operation. Instructions are stored in memory as a sequence of 16-bit little-endian parcels, regardless of memory system endianness. Parcels forming one instruction are stored at increasing halfword addresses, with the lowest-addressed parcel holding the lowest-numbered bits in the instruction specification.

We originally chose little-endian byte ordering for the RISC-V memory system because little-endian systems are currently dominant commercially (all x86 systems; iOS, Android, and Windows for ARM). A minor point is that we have also found little-endian memory systems to be more natural for hardware designers. However, certain application areas, such as IP networking, operate on big-endian data structures, and certain legacy code bases have been built assuming big-endian processors, so we have defined big-endian and bi-endian variants of RISC-V.

We have to fix the order in which instruction parcels are stored in memory, independent of memory system endianness, to ensure that the length-encoding bits always appear first in halfword address order. This allows the length of a variable-length instruction to be quickly determined by an instruction-fetch unit by examining only the first few bits of the first 16-bit instruction parcel.

We further make the instruction parcels themselves little-endian to decouple the instruction encoding from the memory system endianness altogether. This design benefits both software tooling and bi-endian hardware. Otherwise, for instance, a RISC-V assembler or disassembler would always need to know the intended active endianness, despite that in bi-endian systems, the endianness mode might change dynamically during execution. In contrast, by giving instructions a fixed endianness, it is sometimes possible for carefully written software to be endianness-agnostic even in binary form, much like position-independent code.

The choice to have instructions be only little-endian does have consequences, however, for RISC-V software that encodes or decodes machine instructions. Big-endian JIT compilers, for example, must swap the byte order when storing to instruction memory.

Once we had decided to fix on a little-endian instruction encoding, this naturally led to placing the length-encoding bits in the LSB positions of the instruction format to avoid breaking up opcode fields.

1.6. Exceptions, Traps, and Interrupts

We use the term exception to refer to an unusual condition occurring at run time associated with an instruction in the current RISC-V hart. We use the term interrupt to refer to an external asynchronous event that may cause a RISC-V hart to experience an unexpected transfer of control. We use the term trap to refer to the transfer of control to a trap handler caused by either an exception or an interrupt.

The instruction descriptions in following chapters describe conditions that can raise an exception during execution. The general behavior of most RISC-V EEIs is that a trap to some handler occurs when an exception is signaled on an instruction (except for floating-point exceptions, which, in the standard floating-point extensions, do not cause traps). The manner in which interrupts are generated, routed to, and enabled by a hart depends on the EEI.

Our use of "exception" and "trap" is compatible with that in the IEEE-754 floating-point standard.

How traps are handled and made visible to software running on the hart depends on the enclosing execution environment. From the perspective of software running inside an execution environment, traps encountered by a hart at runtime can have four different effects:

Contained Trap

The trap is visible to, and handled by, software running inside the execution environment. For example, in an EEI providing both supervisor and user mode on harts, an ECALL by a user-mode hart will generally result in a transfer of control to a supervisor-mode handler running on the same hart. Similarly, in the same environment, when a hart is interrupted, an interrupt handler will be run in supervisor mode on the hart.

Requested Trap

The trap is a synchronous exception that is an explicit call to the execution environment requesting an action on behalf of software inside the execution environment. An example is a system call. In this case, execution may or may not resume on the hart after the requested action is taken by the execution environment. For example, a system call could remove the hart or cause an orderly termination of the entire execution environment.

Invisible Trap

The trap is handled transparently by the execution environment and execution resumes normally after the trap is handled. Examples include emulating missing instructions, handling non-resident page faults in a demand-paged virtual-memory system, or handling device interrupts for a different job in a multiprogrammed machine. In these cases, the software running inside the execution environment is not aware of the trap (we ignore timing effects in these definitions).

Fatal Trap

The trap represents a fatal failure and causes the execution environment to terminate execution. Examples include failing a virtual-memory page-protection check or allowing a watchdog timer to expire. Each EEI should define how execution is terminated and reported to an external environment.

Table 2 shows the characteristics of each kind of trap.

Table 2. Characteristics of traps
Contained Requested Invisible Fatal

Execution terminates

No

No1

No

Yes

Software is oblivious

No

No

Yes

Yes2

Handled by environment

No

Yes

Yes

Yes

1 Termination may be requested
2 Imprecise fatal traps might be observable by software

The EEI defines for each trap whether it is handled precisely, though the recommendation is to maintain preciseness where possible. Contained and requested traps can be observed to be imprecise by software inside the execution environment. Invisible traps, by definition, cannot be observed to be precise or imprecise by software running inside the execution environment. Fatal traps can be observed to be imprecise by software running inside the execution environment, if known-errorful instructions do not cause immediate termination.

Because this document describes unprivileged instructions, traps are rarely mentioned. Architectural means to handle contained traps are defined in the privileged architecture manual, along with other features to support richer EEIs. Unprivileged instructions that are defined solely to cause requested traps are documented here. Invisible traps are, by their nature, out of scope for this document. Instruction encodings that are not defined here and not defined by some other means may cause a fatal trap.

1.7. UNSPECIFIED Behaviors and Values

The architecture fully describes what implementations must do and any constraints on what they may do. In cases where the architecture intentionally does not constrain implementations, the term UNSPECIFIED is explicitly used.

The term UNSPECIFIED refers to a behavior or value that is intentionally unconstrained. The definition of these behaviors or values is open to extensions, platform standards, or implementations. Extensions, platform standards, or implementation documentation may provide normative content to further constrain cases that the base architecture defines as UNSPECIFIED.

Like the base architecture, extensions should fully describe allowable behavior and values and use the term UNSPECIFIED for cases that are intentionally unconstrained. These cases may be constrained or defined by other extensions, platform standards, or implementations.

2. RV32I Base Integer Instruction Set, Version 2.1

This chapter describes the RV32I base integer instruction set.

RV32I was designed to be sufficient to form a compiler target and to support modern operating system environments. The ISA was also designed to reduce the hardware required in a minimal implementation. RV32I contains 40 unique instructions, though a simple implementation might cover the ECALL/EBREAK instructions with a single SYSTEM hardware instruction that always traps and might be able to implement the FENCE instruction as a NOP, reducing base instruction count to 38 total. RV32I can emulate almost any other ISA extension (except the A extension, which requires additional hardware support for atomicity).

In practice, a hardware implementation including the machine-mode privileged architecture will also require the 6 CSR instructions.

Subsets of the base integer ISA might be useful for pedagogical purposes, but the base has been defined such that there should be little incentive to subset a real hardware implementation beyond omitting support for misaligned memory accesses and treating all SYSTEM instructions as a single trap.

The standard RISC-V assembly language syntax is documented in the Assembly Programmer’s Manual (RISC-V Assembly Programmer’s Manual, n.d.).

Most of the commentary for RV32I also applies to the RV64I base.

2.1. Programmers' Model for Base Integer ISA

Table 3 shows the unprivileged state for the base integer ISA. For RV32I, the 32 x registers are each 32 bits wide, i.e., XLEN=32. Register x0 is hardwired with all bits equal to 0. General purpose registers x1-x31 hold values that various instructions interpret as a collection of Boolean values, or as two’s complement signed binary integers or unsigned binary integers.

There is one additional unprivileged register: the program counter pc holds the address of the current instruction.

Table 3. RISC-V base unprivileged integer register state.
XLEN-1 0

x0/zero

x1

x2

x3

x4

x5

x6

x7

x8

x9

x10

x11

x12

x13

x14

x15

x16

x17

x18

x19

x20

x21

x22

x23

x24

x25

x26

x27

x28

x29

x30

x31

XLEN

XLEN-1

0

pc

XLEN

There is no dedicated stack pointer or subroutine return address link register in the Base Integer ISA; the instruction encoding allows any x register to be used for these purposes. However, the standard software calling convention uses register x1 to hold the return address for a call, with register x5 available as an alternate link register. The standard calling convention uses register x2 as the stack pointer.

Hardware might choose to accelerate function calls and returns that use x1 or x5. See the descriptions of the JAL and JALR instructions.

The optional compressed 16-bit instruction format is designed around the assumption that x1 is the return address register and x2 is the stack pointer. Software using other conventions will operate correctly but may have greater code size.

The number of available architectural registers can have large impacts on code size, performance, and energy consumption. Although 16 registers would arguably be sufficient for an integer ISA running compiled code, it is impossible to encode a complete ISA with 16 registers in 16-bit instructions using a 3-address format. Although a 2-address format would be possible, it would increase instruction count and lower efficiency. We wanted to avoid intermediate instruction sizes (such as Xtensa’s 24-bit instructions) to simplify base hardware implementations, and once a 32-bit instruction size was adopted, it was straightforward to support 32 integer registers. A larger number of integer registers also helps performance on high-performance code, where there can be extensive use of loop unrolling, software pipelining, and cache tiling.

For these reasons, we chose a conventional size of 32 integer registers for RV32I. Dynamic register usage tends to be dominated by a few frequently accessed registers, and regfile implementations can be optimized to reduce access energy for the frequently accessed registers (Tseng & Asanović, 2000). The optional compressed 16-bit instruction format mostly only accesses 8 registers and hence can provide a dense instruction encoding, while additional instruction-set extensions could support a much larger register space (either flat or hierarchical) if desired.

For resource-constrained embedded applications, we have defined the RV32E subset, which only has 16 registers (Chapter 3).

2.2. Base Instruction Formats

In the base RV32I ISA, there are four core instruction formats (R/I/S/U), as shown in Base instruction formats. All are a fixed 32 bits in length. The base ISA has IALIGN=32, meaning that instructions must be aligned on a four-byte boundary in memory. An instruction-address-misaligned exception is generated on a taken branch or unconditional jump if the target address is not IALIGN-bit aligned. This exception is reported on the branch or jump instruction, not on the target instruction. No instruction-address-misaligned exception is generated for a conditional branch that is not taken.

The alignment constraint for base ISA instructions is relaxed to a two-byte boundary when instruction extensions with 16-bit lengths or other odd multiples of 16-bit lengths are added (i.e., IALIGN=16).

Instruction-address-misaligned exceptions are reported on the branch or jump that would cause instruction misalignment to help debugging, and to simplify hardware design for systems with IALIGN=32, where these are the only places where misalignment can occur.

The behavior upon decoding a reserved instruction is UNSPECIFIED.

Some platforms may require that opcodes reserved for standard use raise an illegal-instruction exception. Other platforms may permit reserved opcode space be used for non-conforming extensions.

The RISC-V ISA keeps the source (rs1 and rs2) and destination (rd) registers at the same position in all formats to simplify decoding. Except for the 5-bit immediates used in CSR instructions (Chapter 7), immediates are always sign-extended, and are generally packed towards the leftmost available bits in the instruction and have been allocated to reduce hardware complexity. In particular, the sign bit for all immediates is always in bit 31 of the instruction to speed sign-extension circuitry.

Diagram
Diagram
Diagram
Diagram

RISC-V base instruction formats. Each immediate subfield is labeled with the bit position (imm[x]) in the immediate value being produced, rather than the bit position within the instruction’s immediate field as is usually done.

Decoding register specifiers is usually on the critical paths in implementations, and so the instruction format was chosen to keep all register specifiers at the same position in all formats at the expense of having to move immediate bits across formats (a property shared with RISC-IV aka. SPUR (Lee et al., 1989)).

In practice, most immediates are either small or require all XLEN bits. We chose an asymmetric immediate split (12 bits in regular instructions plus a special load-upper-immediate instruction with 20 bits) to increase the opcode space available for regular instructions.

Immediates are sign-extended because we did not observe a benefit to using zero extension for some immediates as in the MIPS ISA and wanted to keep the ISA as simple as possible.

2.3. Immediate Encoding Variants

There are a further two variants of the instruction formats (B/J) based on the handling of immediates, as shown in Base instruction formats immediate variants..

Diagram
Diagram
Diagram
Diagram
Diagram
Diagram

The only difference between the S and B formats is that the 12-bit immediate field is used to encode branch offsets in multiples of 2 in the B format. Instead of shifting all bits in the instruction-encoded immediate left by one in hardware as is conventionally done, the middle bits (imm[10:1]) and sign bit stay in fixed positions, while the lowest bit in S format (inst[7]) encodes a high-order bit in B format.

Similarly, the only difference between the U and J formats is that the 20-bit immediate is shifted left by 12 bits to form U immediates and by 1 bit to form J immediates. The location of instruction bits in the U and J format immediates is chosen to maximize overlap with the other formats and with each other.

Immediate types shows the immediates produced by each of the base instruction formats, and is labeled to show which instruction bit (inst[y]) produces each bit of the immediate value.

Diagram
Diagram
Diagram
Diagram
Diagram
Figure 1. Types of immediate produced by RISC-V instructions.

The fields are labeled with the instruction bits used to construct their value. Sign extensions always uses inst[31].

Sign extension is one of the most critical operations on immediates (particularly for XLEN>32), and in RISC-V the sign bit for all immediates is always held in bit 31 of the instruction to allow sign extension to proceed in parallel with instruction decoding.

Although more complex implementations might have separate adders for branch and jump calculations and so would not benefit from keeping the location of immediate bits constant across types of instruction, we wanted to reduce the hardware cost of the simplest implementations. By rotating bits in the instruction encoding of B and J immediates instead of using dynamic hardware muxes to multiply the immediate by 2, we reduce instruction signal fanout and immediate mux costs by around a factor of 2. The scrambled immediate encoding will add negligible time to static or ahead-of-time compilation. For dynamic generation of instructions, there is some small additional overhead, but the most common short forward branches have straightforward immediate encodings.

2.4. Integer Computational Instructions

Most integer computational instructions operate on XLEN bits of values held in the integer register file. Integer computational instructions are either encoded as register-immediate operations using the I-type format or as register-register operations using the R-type format. The destination is register rd for both register-immediate and register-register instructions. No integer computational instructions cause arithmetic exceptions.

We did not include special instruction-set support for overflow checks on integer arithmetic operations in the base instruction set, as many overflow checks can be cheaply implemented using RISC-V branches. Overflow checking for unsigned addition requires only a single additional branch instruction after the addition: add t0, t1, t2; bltu t0, t1, overflow.

For signed addition, if one operand’s sign is known, overflow checking requires only a single branch after the addition: addi t0, t1, +imm; blt t0, t1, overflow. This covers the common case of addition with an immediate operand.

For general signed addition, three additional instructions after the addition are required, leveraging the observation that the sum should be less than one of the operands if and only if the other operand is negative.

         add t0, t1, t2
         slti t3, t2, 0
         slt t4, t0, t1
         bne t3, t4, overflow

In RV64I, checks of 32-bit signed additions can be optimized further by comparing the results of ADD and ADDW on the operands.

2.4.1. Integer Register-Immediate Instructions

Diagram

ADDI adds the sign-extended 12-bit immediate to register rs1. Arithmetic overflow is ignored and the result is simply the low XLEN bits of the result. ADDI rd, rs1, 0 is used to implement the MV rd, rs1 assembler pseudoinstruction.

SLTI (set less than immediate) places the value 1 in register rd if register rs1 is less than the sign-extended immediate when both are treated as signed numbers, else 0 is written to rd. SLTIU is similar but compares the values as unsigned numbers (i.e., the immediate is first sign-extended to XLEN bits then treated as an unsigned number). Note, SLTIU rd, rs1, 1 sets rd to 1 if rs1 equals zero, otherwise sets rd to 0 (assembler pseudoinstruction SEQZ rd, rs).

ANDI, ORI, XORI are logical operations that perform bitwise AND, OR, and XOR on register rs1 and the sign-extended 12-bit immediate and place the result in rd. Note, XORI rd, rs1, -1 performs a bitwise logical inversion of register rs1 (assembler pseudoinstruction NOT rd, rs).

Diagram

Shifts by a constant are encoded as a specialization of the I-type format. The operand to be shifted is in rs1, and the shift amount is encoded in the lower 5 bits of the I-immediate field. The right shift type is encoded in bit 30. SLLI is a logical left shift (zeros are shifted into the lower bits); SRLI is a logical right shift (zeros are shifted into the upper bits); and SRAI is an arithmetic right shift (the original sign bit is copied into the vacated upper bits).

Diagram

LUI (load upper immediate) is used to build 32-bit constants and uses the U-type format. LUI places the 32-bit U-immediate value into the destination register rd, filling in the lowest 12 bits with zeros.

AUIPC (add upper immediate to pc) is used to build pc-relative addresses and uses the U-type format. AUIPC forms a 32-bit offset from the U-immediate, filling in the lowest 12 bits with zeros, adds this offset to the address of the AUIPC instruction, then places the result in register rd.

The assembly syntax for lui and auipc does not represent the lower 12 bits of the U-immediate, which are always zero.

The AUIPC instruction supports two-instruction sequences to access arbitrary offsets from the PC for both control-flow transfers and data accesses. The combination of an AUIPC and the 12-bit immediate in a JALR can transfer control to any 32-bit PC-relative address, while an AUIPC plus the 12-bit immediate offset in regular load or store instructions can access any 32-bit PC-relative data address.

The current PC can be obtained by setting the U-immediate to 0. Although a JAL +4 instruction could also be used to obtain the local PC (of the instruction following the JAL), it might cause pipeline breaks in simpler microarchitectures or pollute BTB structures in more complex microarchitectures.

2.4.2. Integer Register-Register Operations

RV32I defines several arithmetic R-type operations. All operations read the rs1 and rs2 registers as source operands and write the result into register rd. The funct7 and funct3 fields select the type of operation.

Diagram

ADD performs the addition of rs1 and rs2. SUB performs the subtraction of rs2 from rs1. Overflows are ignored and the low XLEN bits of results are written to the destination rd. SLT and SLTU perform signed and unsigned compares respectively, writing 1 to rd if rs1 < rs2, 0 otherwise. Note, SLTU rd, x0, rs2 sets rd to 1 if rs2 is not equal to zero, otherwise sets rd to zero (assembler pseudoinstruction SNEZ rd, rs). AND, OR, and XOR perform bitwise logical operations.

SLL, SRL, and SRA perform logical left, logical right, and arithmetic right shifts on the value in register rs1 by the shift amount held in the lower 5 bits of register rs2.

2.4.3. NOP Instruction

Diagram

The NOP instruction does not change any architecturally visible state, except for advancing the pc and incrementing any applicable performance counters. NOP is encoded as ADDI x0, x0, 0.

NOPs can be used to align code segments to microarchitecturally significant address boundaries, or to leave space for inline code modifications. Although there are many possible ways to encode a NOP, we define a canonical NOP encoding to allow microarchitectural optimizations as well as for more readable disassembly output. The other NOP encodings are made available for HINT Instructions.

ADDI was chosen for the NOP encoding as this is most likely to take fewest resources to execute across a range of systems (if not optimized away in decode). In particular, the instruction only reads one register. Also, an ADDI functional unit is more likely to be available in a superscalar design as adds are the most common operation. In particular, address-generation functional units can execute ADDI using the same hardware needed for base+offset address calculations, while register-register ADD or logical/shift operations require additional hardware.

2.5. Control Transfer Instructions

RV32I provides two types of control transfer instructions: unconditional jumps and conditional branches. Control transfer instructions in RV32I do not have architecturally visible delay slots.

If an instruction access-fault or instruction page-fault exception occurs on the target of a jump or taken branch, the exception is reported on the target instruction, not on the jump or branch instruction.

2.5.1. Unconditional Jumps

The jump and link (JAL) instruction uses the J-type format, where the J-immediate encodes a signed offset in multiples of 2 bytes. The offset is sign-extended and added to the address of the jump instruction to form the jump target address. Jumps can therefore target a ±1 MiB range. JAL stores the address of the instruction following the jump ('pc'+4) into register rd. The standard software calling convention uses 'x1' as the return address register and 'x5' as an alternate link register.

The alternate link register supports calling millicode routines (e.g., those to save and restore registers in compressed code) while preserving the regular return address register. The register x5 was chosen as the alternate link register as it maps to a temporary in the standard calling convention, and has an encoding that is only one bit different than the regular link register.

Plain unconditional jumps (assembler pseudoinstruction J) are encoded as a JAL with rd=x0.

Diagram

The indirect jump instruction JALR (jump and link register) uses the I-type encoding. The target address is obtained by adding the sign-extended 12-bit I-immediate to the register rs1, then setting the least-significant bit of the result to zero. The address of the instruction following the jump (pc+4) is written to register rd. Register x0 can be used as the destination if the result is not required.

Plain unconditional indirect jumps (assembler pseudoinstruction JR) are encoded as a JALR with rd=x0. Procedure returns in the standard calling convention (assembler pseudoinstruction RET) are encoded as a JALR with rd=x0, rs1=x1, and imm=0.

Diagram

The unconditional jump instructions all use PC-relative addressing to help support position-independent code. The JALR instruction was defined to enable a two-instruction sequence to jump anywhere in a 32-bit absolute address range. A LUI instruction can first load rs1 with the upper 20 bits of a target address, then JALR can add in the lower bits. Similarly, AUIPC then JALR can jump anywhere in a 32-bit pc-relative address range.

Note that the JALR instruction does not treat the 12-bit immediate as multiples of 2 bytes, unlike the conditional branch instructions. This avoids one more immediate format in hardware. In practice, most uses of JALR will have either a zero immediate or be paired with a LUI or AUIPC, so the slight reduction in range is not significant.

Clearing the least-significant bit when calculating the JALR target address both simplifies the hardware slightly and allows the low bit of function pointers to be used to store auxiliary information. Although there is potentially a slight loss of error checking in this case, in practice jumps to an incorrect instruction address will usually quickly raise an exception.

When used with a base rs1=x0, JALR can be used to implement a single instruction subroutine call to the lowest or highest address region from anywhere in the address space, which could be used to implement fast calls to a small runtime library. Alternatively, an ABI could dedicate a general-purpose register to point to a library elsewhere in the address space.

The JAL and JALR instructions will generate an instruction-address-misaligned exception if the target address is not aligned to a four-byte boundary.

Instruction-address-misaligned exceptions are not possible on machines that support extensions with 16-bit aligned instructions, such as the compressed instruction-set extension, C.

Return-address prediction stacks are a common feature of high-performance instruction-fetch units, but require accurate detection of instructions used for procedure calls and returns to be effective. For RISC-V, hints as to the instructions' usage are encoded implicitly via the register numbers used. A JAL instruction should push the return address onto a return-address stack (RAS) only when rd is 'x1' or x5. JALR instructions should push/pop a RAS as shown in Table 4.

Table 4. Return-address stack prediction hints encoded in the register operands of a JALR instruction.
rd is x1/x5 rs1 is x1/x5 rd=rs1 RAS action

No

No

 — 

None

No

Yes

 — 

Pop

Yes

No

 — 

Push

Yes

Yes

No

Pop, then push

Yes

Yes

Yes

Push

Some other ISAs added explicit hint bits to their indirect-jump instructions to guide return-address stack manipulation. We use implicit hinting tied to register numbers and the calling convention to reduce the encoding space used for these hints.

When two different link registers (x1 and x5) are given as rs1 and rd, then the RAS is both popped and pushed to support coroutines. If rs1 and rd are the same link register (either x1 or x5), the RAS is only pushed to enable macro-op fusion of the sequences: lui ra, imm20; jalr ra, imm12(ra)_ and _auipc ra, imm20; jalr ra, imm12(ra)

2.5.2. Conditional Branches

All branch instructions use the B-type instruction format. The 12-bit B-immediate encodes signed offsets in multiples of 2 bytes. The offset is sign-extended and added to the address of the branch instruction to give the target address. The conditional branch range is ±4 KiB.

Diagram

Branch instructions compare two registers. BEQ and BNE take the branch if registers rs1 and rs2 are equal or unequal respectively. BLT and BLTU take the branch if rs1 is less than rs2, using signed and unsigned comparison respectively. BGE and BGEU take the branch if rs1 is greater than or equal to rs2, using signed and unsigned comparison respectively. Note, BGT, BGTU, BLE, and BLEU can be synthesized by reversing the operands to BLT, BLTU, BGE, and BGEU, respectively.

Signed array bounds may be checked with a single BLTU instruction, since any negative index will compare greater than any nonnegative bound.

Software should be optimized such that the sequential code path is the most common path, with less-frequently taken code paths placed out of line. Software should also assume that backward branches will be predicted taken and forward branches as not taken, at least the first time they are encountered. Dynamic predictors should quickly learn any predictable branch behavior.

Unlike some other architectures, the RISC-V jump (JAL with rd=x0) instruction should always be used for unconditional branches instead of a conditional branch instruction with an always-true condition. RISC-V jumps are also PC-relative and support a much wider offset range than branches, and will not pollute conditional-branch prediction tables.

The conditional branches were designed to include arithmetic comparison operations between two registers (as also done in PA-RISC, Xtensa, and MIPS R6), rather than use condition codes (x86, ARM, SPARC, PowerPC), or to only compare one register against zero (Alpha, MIPS), or two registers only for equality (MIPS). This design was motivated by the observation that a combined compare-and-branch instruction fits into a regular pipeline, avoids additional condition code state or use of a temporary register, and reduces static code size and dynamic instruction fetch traffic. Another point is that comparisons against zero require non-trivial circuit delay (especially after the move to static logic in advanced processes) and so are almost as expensive as arithmetic magnitude compares. Another advantage of a fused compare-and-branch instruction is that branches are observed earlier in the front-end instruction stream, and so can be predicted earlier. There is perhaps an advantage to a design with condition codes in the case where multiple branches can be taken based on the same condition codes, but we believe this case to be relatively rare.

We considered but did not include static branch hints in the instruction encoding. These can reduce the pressure on dynamic predictors, but require more instruction encoding space and software profiling for best results, and can result in poor performance if production runs do not match profiling runs.

We considered but did not include conditional moves or predicated instructions, which can effectively replace unpredictable short forward branches. Conditional moves are the simpler of the two, but are difficult to use with conditional code that might cause exceptions (memory accesses and floating-point operations). Predication adds additional flag state to a system, additional instructions to set and clear flags, and additional encoding overhead on every instruction. Both conditional move and predicated instructions add complexity to out-of-order microarchitectures, adding an implicit third source operand due to the need to copy the original value of the destination architectural register into the renamed destination physical register if the predicate is false. Also, static compile-time decisions to use predication instead of branches can result in lower performance on inputs not included in the compiler training set, especially given that unpredictable branches are rare, and becoming rarer as branch prediction techniques improve.

We note that various microarchitectural techniques exist to dynamically convert unpredictable short forward branches into internally predicated code to avoid the cost of flushing pipelines on a branch mispredict (Heil & Smith, 1996), (Klauser et al., 1998), (Kim et al., 2005) and have been implemented in commercial processors (Sinharoy et al., 2011). The simplest techniques just reduce the penalty of recovering from a mispredicted short forward branch by only flushing instructions in the branch shadow instead of the entire fetch pipeline, or by fetching instructions from both sides using wide instruction fetch or idle instruction fetch slots. More complex techniques for out-of-order cores add internal predicates on instructions in the branch shadow, with the internal predicate value written by the branch instruction, allowing the branch and following instructions to be executed speculatively and out-of-order with respect to other code.

The conditional branch instructions will generate an instruction-address-misaligned exception if the target address is not aligned to a four-byte boundary and the branch condition evaluates to true. If the branch condition evaluates to false, the instruction-address-misaligned exception will not be raised.

Instruction-address-misaligned exceptions are not possible on machines that support extensions with 16-bit aligned instructions, such as the compressed instruction-set extension, C.

2.6. Load and Store Instructions

RV32I is a load-store architecture, where only load and store instructions access memory and arithmetic instructions only operate on CPU registers. RV32I provides a 32-bit address space that is byte-addressed. The EEI will define what portions of the address space are legal to access with which instructions (e.g., some addresses might be read only, or support word access only). Loads with a destination of x0 must still raise any exceptions and cause any other side effects even though the load value is discarded.

The EEI will define whether the memory system is little-endian or big-endian. In RISC-V, endianness is byte-address invariant.

In a system for which endianness is byte-address invariant, the following property holds: if a byte is stored to memory at some address in some endianness, then a byte-sized load from that address in any endianness returns the stored value.

In a little-endian configuration, multibyte stores write the least-significant register byte at the lowest memory byte address, followed by the other register bytes in ascending order of their significance. Loads similarly transfer the contents of the lesser memory byte addresses to the less-significant register bytes.

In a big-endian configuration, multibyte stores write the most-significant register byte at the lowest memory byte address, followed by the other register bytes in descending order of their significance. Loads similarly transfer the contents of the greater memory byte addresses to the less-significant register bytes.
Diagram
Diagram

Load and store instructions transfer a value between the registers and memory. Loads are encoded in the I-type format and stores are S-type. The effective address is obtained by adding register rs1 to the sign-extended 12-bit offset. Loads copy a value from memory to register rd. Stores copy the value in register rs2 to memory.

The LW instruction loads a 32-bit value from memory into rd. LH loads a 16-bit value from memory, then sign-extends to 32-bits before storing in rd. LHU loads a 16-bit value from memory but then zero extends to 32-bits before storing in rd. LB and LBU are defined analogously for 8-bit values. The SW, SH, and SB instructions store 32-bit, 16-bit, and 8-bit values from the low bits of register rs2 to memory.

Regardless of EEI, loads and stores whose effective addresses are naturally aligned shall not raise an address-misaligned exception. Loads and stores whose effective address is not naturally aligned to the referenced datatype (i.e., the effective address is not divisible by the size of the access in bytes) have behavior dependent on the EEI.

An EEI may guarantee that misaligned loads and stores are fully supported, and so the software running inside the execution environment will never experience a contained or fatal address-misaligned trap. In this case, the misaligned loads and stores can be handled in hardware, or via an invisible trap into the execution environment implementation, or possibly a combination of hardware and invisible trap depending on address.

An EEI may not guarantee misaligned loads and stores are handled invisibly. In this case, loads and stores that are not naturally aligned may either complete execution successfully or raise an exception. The exception raised can be either an address-misaligned exception or an access-fault exception. For a memory access that would otherwise be able to complete except for the misalignment, an access-fault exception can be raised instead of an address-misaligned exception if the misaligned access should not be emulated, e.g., if accesses to the memory region have side effects. When an EEI does not guarantee misaligned loads and stores are handled invisibly, the EEI must define if exceptions caused by address misalignment result in a contained trap (allowing software running inside the execution environment to handle the trap) or a fatal trap (terminating execution).

Misaligned accesses are occasionally required when porting legacy code, and help performance on applications when using any form of packed-SIMD extension or handling externally packed data structures. Our rationale for allowing EEIs to choose to support misaligned accesses via the regular load and store instructions is to simplify the addition of misaligned hardware support. One option would have been to disallow misaligned accesses in the base ISAs and then provide some separate ISA support for misaligned accesses, either special instructions to help software handle misaligned accesses or a new hardware addressing mode for misaligned accesses. Special instructions are difficult to use, complicate the ISA, and often add new processor state (e.g., SPARC VIS align address offset register) or complicate access to existing processor state (e.g., MIPS LWL/LWR partial register writes). In addition, for loop-oriented packed-SIMD code, the extra overhead when operands are misaligned motivates software to provide multiple forms of loop depending on operand alignment, which complicates code generation and adds to loop startup overhead. New misaligned hardware addressing modes take considerable space in the instruction encoding or require very simplified addressing modes (e.g., register indirect only).

Even when misaligned loads and stores complete successfully, these accesses might run extremely slowly depending on the implementation (e.g., when implemented via an invisible trap). Furthermore, whereas naturally aligned loads and stores are guaranteed to execute atomically, misaligned loads and stores might not, and hence require additional synchronization to ensure atomicity.

We do not mandate atomicity for misaligned accesses so execution environment implementations can use an invisible machine trap and a software handler to handle some or all misaligned accesses. If hardware misaligned support is provided, software can exploit this by simply using regular load and store instructions. Hardware can then automatically optimize accesses depending on whether runtime addresses are aligned.

2.7. Memory Ordering Instructions

mem order

The FENCE instruction is used to order device I/O and memory accesses as viewed by other RISC-V harts and external devices or coprocessors. Any combination of device input (I), device output (O), memory reads (R), and memory writes (W) may be ordered with respect to any combination of the same. Informally, no other RISC-V hart or external device can observe any operation in the successor set following a FENCE before any operation in the predecessor set preceding the FENCE. Chapter 18 provides a precise description of the RISC-V memory consistency model.

The FENCE instruction also orders memory reads and writes made by the hart as observed by memory reads and writes made by an external device. However, FENCE does not order observations of events made by an external device using any other signaling mechanism.

A device might observe an access to a memory location via some external communication mechanism, e.g., a memory-mapped control register that drives an interrupt signal to an interrupt controller. This communication is outside the scope of the FENCE ordering mechanism and hence the FENCE instruction can provide no guarantee on when a change in the interrupt signal is visible to the interrupt controller. Specific devices might provide additional ordering guarantees to reduce software overhead but those are outside the scope of the RISC-V memory model.

The EEI will define what I/O operations are possible, and in particular, which memory addresses when accessed by load and store instructions will be treated and ordered as device input and device output operations respectively rather than memory reads and writes. For example, memory-mapped I/O devices will typically be accessed with uncached loads and stores that are ordered using the I and O bits rather than the R and W bits. Instruction-set extensions might also describe new I/O instructions that will also be ordered using the I and O bits in a FENCE.

Table 5. Fence mode encoding
fm field Mnemonic Meaning

0000

none

Normal Fence

1000

TSO

With FENCE RW,RW: exclude write-to-read ordering; otherwise: Reserved for future use.

other

Reserved for future use.

The fence mode field fm defines the semantics of the FENCE. A FENCE with fm=0000 orders all memory operations in its predecessor set before all memory operations in its successor set.

The FENCE.TSO instruction is encoded as a FENCE instruction with fm=1000, predecessor=RW, and successor=RW. FENCE.TSO orders all load operations in its predecessor set before all memory operations in its successor set, and all store operations in its predecessor set before all store operations in its successor set. This leaves non-AMO store operations in the FENCE.TSO’s predecessor set unordered with non-AMO loads in its successor set.

Because FENCE RW,RW imposes a superset of the orderings that FENCE.TSO imposes, it is correct to ignore the fm field and implement FENCE.TSO as FENCE RW,RW.

The unused fields in the FENCE instructions--rs1 and rd--are reserved for finer-grain fences in future extensions. For forward compatibility, base implementations shall ignore these fields, and standard software shall zero these fields. Likewise, many fm and predecessor/successor set settings in Table 5 are also reserved for future use. Base implementations shall treat all such reserved configurations as normal fences with fm=0000, and standard software shall use only non-reserved configurations.

We chose a relaxed memory model to allow high performance from simple machine implementations and from likely future coprocessor or accelerator extensions. We separate out I/O ordering from memory R/W ordering to avoid unnecessary serialization within a device-driver hart and also to support alternative non-memory paths to control added coprocessors or I/O devices. Simple implementations may additionally ignore the predecessor and successor fields and always execute a conservative fence on all operations.

2.8. Environment Call and Breakpoints

SYSTEM instructions are used to access system functionality that might require privileged access and are encoded using the I-type instruction format. These can be divided into two main classes: those that atomically read-modify-write control and status registers (CSRs), and all other potentially privileged instructions. CSR instructions are described in Chapter 7, and the base unprivileged instructions are described in the following section.

The SYSTEM instructions are defined to allow simpler implementations to always trap to a single software trap handler. More sophisticated implementations might execute more of each system instruction in hardware.
Diagram

These two instructions cause a precise requested trap to the supporting execution environment.

The ECALL instruction is used to make a service request to the execution environment. The EEI will define how parameters for the service request are passed, but usually these will be in defined locations in the integer register file.

The EBREAK instruction is used to return control to a debugging environment.

ECALL and EBREAK were previously named SCALL and SBREAK. The instructions have the same functionality and encoding, but were renamed to reflect that they can be used more generally than to call a supervisor-level operating system or debugger.

EBREAK was primarily designed to be used by a debugger to cause execution to stop and fall back into the debugger. EBREAK is also used by the standard gcc compiler to mark code paths that should not be executed.

Another use of EBREAK is to support "semihosting", where the execution environment includes a debugger that can provide services over an alternate system call interface built around the EBREAK instruction. Because the RISC-V base ISAs do not provide more than one EBREAK instruction, RISC-V semihosting uses a special sequence of instructions to distinguish a semihosting EBREAK from a debugger inserted EBREAK.

    slli x0, x0, 0x1f   # Entry NOP
    ebreak              # Break to debugger
    srai x0, x0, 7      # Exit NOP

Note that these three instructions must be 32-bit-wide instructions, i.e., they mustn’t be among the compressed 16-bit instructions described in Chapter 28.

The shift NOP instructions are still considered available for use as HINTs.

Semihosting is a form of service call and would be more naturally encoded as an ECALL using an existing ABI, but this would require the debugger to be able to intercept ECALLs, which is a newer addition to the debug standard. We intend to move over to using ECALLs with a standard ABI, in which case, semihosting can share a service ABI with an existing standard.

We note that ARM processors have also moved to using SVC instead of BKPT for semihosting calls in newer designs.

2.9. HINT Instructions

RV32I reserves a large encoding space for HINT instructions, which are usually used to communicate performance hints to the microarchitecture. Like the NOP instruction, HINTs do not change any architecturally visible state, except for advancing the pc and any applicable performance counters. Implementations are always allowed to ignore the encoded hints.

Most RV32I HINTs are encoded as integer computational instructions with rd=x0. The other RV32I HINTs are encoded as FENCE instructions with a null predecessor or successor set and with fm=0.

These HINT encodings have been chosen so that simple implementations can ignore HINTs altogether, and instead execute a HINT as a regular instruction that happens not to mutate the architectural state. For example, ADD is a HINT if the destination register is x0; the five-bit rs1 and rs2 fields encode arguments to the HINT. However, a simple implementation can simply execute the HINT as an ADD of rs1 and rs2 that writes x0, which has no architecturally visible effect.

As another example, a FENCE instruction with a zero pred field and a zero fm field is a HINT; the succ, rs1, and rd fields encode the arguments to the HINT. A simple implementation can simply execute the HINT as a FENCE that orders the null set of prior memory accesses before whichever subsequent memory accesses are encoded in the succ field. Since the intersection of the predecessor and successor sets is null, the instruction imposes no memory orderings, and so it has no architecturally visible effect.

Table 6 lists all RV32I HINT code points. 91% of the HINT space is reserved for standard HINTs. The remainder of the HINT space is designated for custom HINTs: no standard HINTs will ever be defined in this subspace.

We anticipate standard hints to eventually include memory-system spatial and temporal locality hints, branch prediction hints, thread-scheduling hints, security tags, and instrumentation flags for simulation/emulation.
Table 6. RV32I HINT instructions.
Instruction Constraints Code Points Purpose

LUI

rd=x0

stem 3d91b2575aebc0d742415f31a7138964

Designated for future standard use

AUIPC

rd=x0

stem 3d91b2575aebc0d742415f31a7138964

ADDI

rd=x0, and either rs1≠x0 or imm≠0

stem 0f73b0acd2f2b2250dca11c3c9ec1713

ANDI

rd=x0

stem cec3b648d9a79b7951e2288a4cb712cd

ORI

rd=x0

stem cec3b648d9a79b7951e2288a4cb712cd

XORI

rd=x0

stem cec3b648d9a79b7951e2288a4cb712cd

ADD

rd=x0, rs1≠x0

stem e68410fd305668d5d27741524c852b69

ADD

rd=x0, rs1=x0, rs2≠x2-x5

28

ADD

rd=x0, rs1=x0, rs2=x2-x5

4

(rs2=x2) NTL.P1
(rs2=x3) NTL.PALL
(rs2=x4) NTL.S1
(rs2=x5) NTL.ALL

SUB

rd=x0

stem 560d002b3cbfbdc26102752336160ae5

Designated for future standard use

AND

rd=x0

stem 560d002b3cbfbdc26102752336160ae5

OR

rd=x0

stem 560d002b3cbfbdc26102752336160ae5

XOR

rd=x0

stem 560d002b3cbfbdc26102752336160ae5

SLL

rd=x0

stem 560d002b3cbfbdc26102752336160ae5

SRL

rd=x0

stem 560d002b3cbfbdc26102752336160ae5

SRA

rd=x0

stem 560d002b3cbfbdc26102752336160ae5

FENCE

rd=x0, rs1≠x0, fm=0, and either pred=0 or succ=0

stem 3be453f8a29ef791f6184cf706174789

FENCE

rd≠x0, rs1=x0, fm=0, and either pred=0 or succ=0

stem 3be453f8a29ef791f6184cf706174789

FENCE

rd=rs1=x0, fm=0, pred=0, succ≠0

15

FENCE

rd=rs1=x0, fm=0, pred≠W, succ=0

15

FENCE

rd=rs1=x0, fm=0, pred=W, succ=0

1

PAUSE

SLTI

rd=x0

stem cec3b648d9a79b7951e2288a4cb712cd

Designated for custom use

SLTIU

rd=x0

stem cec3b648d9a79b7951e2288a4cb712cd

SLLI

rd=x0

stem 560d002b3cbfbdc26102752336160ae5

SRLI

rd=x0

stem 560d002b3cbfbdc26102752336160ae5

SRAI

rd=x0

stem 560d002b3cbfbdc26102752336160ae5

SLT

rd=x0

stem 560d002b3cbfbdc26102752336160ae5

SLTU

rd=x0

stem 560d002b3cbfbdc26102752336160ae5
When allocating slli x0, x0, 0x1f or srai x0, x0, 7 as custom HINTs, take note of their use in semihosting calls, as described in Section 2.8.

3. RV32E and RV64E Base Integer Instruction Sets, Version 2.0

CV32A65X: These instruction sets are not supported.

4. RV64I Base Integer Instruction Set, Version 2.1

CV32A65X: This instruction set is not supported.

5. RV128I Base Integer Instruction Set, Version 1.7

CV32A65X: This instruction set is not supported.

6. "Zifencei" Extension for Instruction-Fetch Fence, Version 2.0

CV32A65X: This extension is not supported.

7. "Zicsr", Extension for Control and Status Register (CSR) Instructions, Version 2.0

RISC-V defines a separate address space of 4096 Control and Status registers associated with each hart. This chapter defines the full set of CSR instructions that operate on these CSRs.

While CSRs are primarily used by the privileged architecture, there are several uses in unprivileged code including for counters and timers, and for floating-point status.

The counters and timers are no longer considered mandatory parts of the standard base ISAs, and so the CSR instructions required to access them have been moved out of Chapter 2 into this separate chapter.

7.1. CSR Instructions

All CSR instructions atomically read-modify-write a single CSR, whose CSR specifier is encoded in the 12-bit csr field of the instruction held in bits 31-20. The immediate forms use a 5-bit zero-extended immediate encoded in the rs1 field.

Diagram

The CSRRW (Atomic Read/Write CSR) instruction atomically swaps values in the CSRs and integer registers. CSRRW reads the old value of the CSR, zero-extends the value to XLEN bits, then writes it to integer register rd. The initial value in rs1 is written to the CSR. If rd=x0, then the instruction shall not read the CSR and shall not cause any of the side effects that might occur on a CSR read.

The CSRRS (Atomic Read and Set Bits in CSR) instruction reads the value of the CSR, zero-extends the value to XLEN bits, and writes it to integer register rd. The initial value in integer register rs1 is treated as a bit mask that specifies bit positions to be set in the CSR. Any bit that is high in rs1 will cause the corresponding bit to be set in the CSR, if that CSR bit is writable.

The CSRRC (Atomic Read and Clear Bits in CSR) instruction reads the value of the CSR, zero-extends the value to XLEN bits, and writes it to integer register rd. The initial value in integer register rs1 is treated as a bit mask that specifies bit positions to be cleared in the CSR. Any bit that is high in rs1 will cause the corresponding bit to be cleared in the CSR, if that CSR bit is writable.

For both CSRRS and CSRRC, if rs1=x0, then the instruction will not write to the CSR at all, and so shall not cause any of the side effects that might otherwise occur on a CSR write, nor raise illegal-instruction exceptions on accesses to read-only CSRs. Both CSRRS and CSRRC always read the addressed CSR and cause any read side effects regardless of rs1 and rd fields. Note that if rs1 specifies a register other than x0, and that register holds a zero value, the instruction will not action any attendant per-field side effects, but will action any side effects caused by writing to the entire CSR.

A CSRRW with rs1=x0 will attempt to write zero to the destination CSR.

The CSRRWI, CSRRSI, and CSRRCI variants are similar to CSRRW, CSRRS, and CSRRC respectively, except they update the CSR using an XLEN-bit value obtained by zero-extending a 5-bit unsigned immediate (uimm[4:0]) field encoded in the rs1 field instead of a value from an integer register. For CSRRSI and CSRRCI, if the uimm[4:0] field is zero, then these instructions will not write to the CSR, and shall not cause any of the side effects that might otherwise occur on a CSR write, nor raise illegal-instruction exceptions on accesses to read-only CSRs. For CSRRWI, if rd=x0, then the instruction shall not read the CSR and shall not cause any of the side effects that might occur on a CSR read. Both CSRRSI and CSRRCI will always read the CSR and cause any read side effects regardless of rd and rs1 fields.

Table 7. Conditions determining whether a CSR instruction reads or writes the specified CSR.
Register operand

Instruction

rd is x0

rs1 is x0

Reads CSR

Writes CSR

CSRRW

Yes

-

No

Yes

CSRRW

No

-

Yes

Yes

CSRRS/CSRRC

-

Yes

Yes

No

CSRRS/CSRRC

-

No

Yes

Yes

Immediate operand

Instruction

rd is x0

uimmstem bf536b3e7d45040baab197f00bea9eba0

Reads CSR

Writes CSR

CSRRWI

Yes

-

No

Yes

CSRRWI

No

-

Yes

Yes

CSRRSI/CSRRCI

-

Yes

Yes

No

CSRRSI/CSRRCI

-

No

Yes

Yes

Table 7 summarizes the behavior of the CSR instructions with respect to whether they read and/or write the CSR.

In addition to side effects that occur as a consequence of reading or writing a CSR, individual fields within a CSR might have side effects when written. The CSRRW[I] instructions action side effects for all such fields within the written CSR. The CSRRS[I] an CSRRC[I] instructions only action side effects for fields for which the rs1 or uimm argument has at least one bit set corresponding to that field.

As of this writing, no standard CSRs have side effects on field writes. Hence, whether a standard CSR access has any side effects can be determined solely from the opcode.

Defining CSRs with side effects on field writes is not recommended.

For any event or consequence that occurs due to a CSR having a particular value, if a write to the CSR gives it that value, the resulting event or consequence is said to be an indirect effect of the write. Indirect effects of a CSR write are not considered by the RISC-V ISA to be side effects of that write.

An example of side effects for CSR accesses would be if reading from a specific CSR causes a light bulb to turn on, while writing an odd value to the same CSR causes the light to turn off. Assume writing an even value has no effect. In this case, both the read and write have side effects controlling whether the bulb is lit, as this condition is not determined solely from the CSR value. (Note that after writing an odd value to the CSR to turn off the light, then reading to turn the light on, writing again the same odd value causes the light to turn off again. Hence, on the last write, it is not a change in the CSR value that turns off the light.)

On the other hand, if a bulb is rigged to light whenever the value of a particular CSR is odd, then turning the light on and off is not considered a side effect of writing to the CSR but merely an indirect effect of such writes.

More concretely, the RISC-V privileged architecture defined in Volume II specifies that certain combinations of CSR values cause a trap to occur. When an explicit write to a CSR creates the conditions that trigger the trap, the trap is not considered a side effect of the write but merely an indirect effect.

Standard CSRs do not have any side effects on reads. Standard CSRs may have side effects on writes. Custom extensions might add CSRs for which accesses have side effects on either reads or writes.

Some CSRs, such as the instructions-retired counter, instret, may be modified as side effects of instruction execution. In these cases, if a CSR access instruction reads a CSR, it reads the value prior to the execution of the instruction. If a CSR access instruction writes such a CSR, the explicit write is done instead of the update from the side effect. In particular, a value written to instret by one instruction will be the value read by the following instruction.

The assembler pseudoinstruction to read a CSR, CSRR rd, csr, is encoded as CSRRS rd, csr, x0. The assembler pseudoinstruction to write a CSR, CSRW csr, rs1, is encoded as CSRRW x0, csr, rs1, while CSRWI csr, uimm, is encoded as CSRRWI x0, csr, uimm.

Further assembler pseudoinstructions are defined to set and clear bits in the CSR when the old value is not required: CSRS/CSRC csr, rs1; CSRSI/CSRCI csr, uimm.

7.1.1. CSR Access Ordering

Each RISC-V hart normally observes its own CSR accesses, including its implicit CSR accesses, as performed in program order. In particular, unless specified otherwise, a CSR access is performed after the execution of any prior instructions in program order whose behavior modifies or is modified by the CSR state and before the execution of any subsequent instructions in program order whose behavior modifies or is modified by the CSR state. Furthermore, an explicit CSR read returns the CSR state before the execution of the instruction, while an explicit CSR write suppresses and overrides any implicit writes or modifications to the same CSR by the same instruction.

Likewise, any side effects from an explicit CSR access are normally observed to occur synchronously in program order. Unless specified otherwise, the full consequences of any such side effects are observable by the very next instruction, and no consequences may be observed out-of-order by preceding instructions. (Note the distinction made earlier between side effects and indirect effects of CSR writes.)

For the RVWMO memory consistency model (Chapter 18), CSR accesses are weakly ordered by default, so other harts or devices may observe CSR accesses in an order different from program order. In addition, CSR accesses are not ordered with respect to explicit memory accesses, unless a CSR access modifies the execution behavior of the instruction that performs the explicit memory access or unless a CSR access and an explicit memory access are ordered by either the syntactic dependencies defined by the memory model or the ordering requirements defined by the Memory-Ordering PMAs section in Volume II of this manual. To enforce ordering in all other cases, software should execute a FENCE instruction between the relevant accesses. For the purposes of the FENCE instruction, CSR read accesses are classified as device input (I), and CSR write accesses are classified as device output (O).

Informally, the CSR space acts as a weakly ordered memory-mapped I/O region, as defined by the Memory-Ordering PMAs section in Volume II of this manual. As a result, the order of CSR accesses with respect to all other accesses is constrained by the same mechanisms that constrain the order of memory-mapped I/O accesses to such a region.

These CSR-ordering constraints are imposed to support ordering main memory and memory-mapped I/O accesses with respect to CSR accesses that are visible to, or affected by, devices or other harts. Examples include the time, cycle, and mcycle CSRs, in addition to CSRs that reflect pending interrupts, like mip and sip. Note that implicit reads of such CSRs (e.g., taking an interrupt because of a change in mip) are also ordered as device input.

Most CSRs (including, e.g., the fcsr) are not visible to other harts; their accesses can be freely reordered in the global memory order with respect to FENCE instructions without violating this specification.

The hardware platform may define that accesses to certain CSRs are strongly ordered, as defined by the Memory-Ordering PMAs section in Volume II of this manual. Accesses to strongly ordered CSRs have stronger ordering constraints with respect to accesses to both weakly ordered CSRs and accesses to memory-mapped I/O regions.

The rules for the reordering of CSR accesses in the global memory order should probably be moved to Chapter 18 concerning the RVWMO memory consistency model.

8. "Zicntr" and "Zihpm" Extensions for Counters, Version 2.0

RISC-V ISAs provide a set of up to thirty-two 64-bit performance counters and timers that are accessible via unprivileged XLEN-bit read-only CSR registers 0xC00–0xC1F (when XLEN=32, the upper 32 bits are accessed via CSR registers 0xC80–0xC9F). These counters are divided between the "Zicntr" and "Zihpm" extensions.

8.1. "Zicntr" Extension for Base Counters and Timers

The Zicntr standard extension comprises the first three of these counters (CYCLE, TIME, and INSTRET), which have dedicated functions (cycle count, real-time clock, and instructions retired, respectively). The Zicntr extension depends on the Zicsr extension.

We recommend provision of these basic counters in implementations as they are essential for basic performance analysis, adaptive and dynamic optimization, and to allow an application to work with real-time streams. Additional counters in the separate Zihpm extension can help diagnose performance problems and these should be made accessible from user-level application code with low overhead.

Some execution environments might prohibit access to counters, for example, to impede timing side-channel attacks.
Diagram

For base ISAs with XLEN≥64, CSR instructions can access the full 64-bit CSRs directly. In particular, the RDCYCLE, RDTIME, and RDINSTRET pseudoinstructions read the full 64 bits of the cycle, time, and instret counters.

The counter pseudoinstructions are mapped to the read-only csrrs rd, counter, x0 canonical form, but the other read-only CSR instruction forms (based on CSRRC/CSRRSI/CSRRCI) are also legal ways to read these CSRs.

For base ISAs with XLEN=32, the Zicntr extension enables the three 64-bit read-only counters to be accessed in 32-bit pieces. The RDCYCLE, RDTIME, and RDINSTRET pseudoinstructions provide the lower 32 bits, and the RDCYCLEH, RDTIMEH, and RDINSTRETH pseudoinstructions provide the upper 32 bits of the respective counters.

We required the counters be 64 bits wide, even when XLEN=32, as otherwise it is very difficult for software to determine if values have overflowed. For a low-end implementation, the upper 32 bits of each counter can be implemented using software counters incremented by a trap handler triggered by overflow of the lower 32 bits. The sample code given below shows how the full 64-bit width value can be safely read using the individual 32-bit width pseudoinstructions.

The RDCYCLE pseudoinstruction reads the low XLEN bits of the cycle CSR which holds a count of the number of clock cycles executed by the processor core on which the hart is running from an arbitrary start time in the past. RDCYCLEH is only present when XLEN=32 and reads bits 63-32 of the same cycle counter. The underlying 64-bit counter should never overflow in practice. The rate at which the cycle counter advances will depend on the implementation and operating environment. The execution environment should provide a means to determine the current rate (cycles/second) at which the cycle counter is incrementing.

RDCYCLE is intended to return the number of cycles executed by the processor core, not the hart. Precisely defining what is a "core" is difficult given some implementation choices (e.g., AMD Bulldozer). Precisely defining what is a "clock cycle" is also difficult given the range of implementations (including software emulations), but the intent is that RDCYCLE is used for performance monitoring along with the other performance counters. In particular, where there is one hart/core, one would expect cycle-count/instructions-retired to measure CPI for a hart.

Cores don’t have to be exposed to software at all, and an implementor might choose to pretend multiple harts on one physical core are running on separate cores with one hart/core, and provide separate cycle counters for each hart. This might make sense in a simple barrel processor (e.g., CDC 6600 peripheral processors) where inter-hart timing interactions are non-existent or minimal.

Where there is more than one hart/core and dynamic multithreading, it is not generally possible to separate out cycles per hart (especially with SMT). It might be possible to define a separate performance counter that tried to capture the number of cycles a particular hart was running, but this definition would have to be very fuzzy to cover all the possible threading implementations. For example, should we only count cycles for which any instruction was issued to execution for this hart, and/or cycles any instruction retired, or include cycles this hart was occupying machine resources but couldn’t execute due to stalls while other harts went into execution? Likely, "all of the above" would be needed to have understandable performance stats. This complexity of defining a per-hart cycle count, and also the need in any case for a total per-core cycle count when tuning multithreaded code led to just standardizing the per-core cycle counter, which also happens to work well for the common single hart/core case.

Standardizing what happens during "sleep" is not practical given that what "sleep" means is not standardized across execution environments, but if the entire core is paused (entirely clock-gated or powered-down in deep sleep), then it is not executing clock cycles, and the cycle count shouldn’t be increasing per the spec. There are many details, e.g., whether clock cycles required to reset a processor after waking up from a power-down event should be counted, and these are considered execution-environment-specific details.

Even though there is no precise definition that works for all platforms, this is still a useful facility for most platforms, and an imprecise, common, "usually correct" standard here is better than no standard. The intent of RDCYCLE was primarily performance monitoring/tuning, and the specification was written with that goal in mind.

The RDTIME pseudoinstruction reads the low XLEN bits of the "time" CSR, which counts wall-clock real time that has passed from an arbitrary start time in the past. RDTIMEH is only present when XLEN=32 and reads bits 63-32 of the same real-time counter. The underlying 64-bit counter increments by one with each tick of the real-time clock, and, for realistic real-time clock frequencies, should never overflow in practice. The execution environment should provide a means of determining the period of a counter tick (seconds/tick). The period should be constant within a small error bound. The environment should provide a means to determine the accuracy of the clock (i.e., the maximum relative error between the nominal and actual real-time clock periods).

On some simple platforms, cycle count might represent a valid implementation of RDTIME, in which case RDTIME and RDCYCLE may return the same result.

It is difficult to provide a strict mandate on clock period given the wide variety of possible implementation platforms. The maximum error bound should be set based on the requirements of the platform.

The real-time clocks of all harts must be synchronized to within one tick of the real-time clock.

As with other architectural mandates, it suffices to appear "as if" harts are synchronized to within one tick of the real-time clock, i.e., software is unable to observe that there is a greater delta between the real-time clock values observed on two harts.

The RDINSTRET pseudoinstruction reads the low XLEN bits of the instret CSR, which counts the number of instructions retired by this hart from some arbitrary start point in the past. RDINSTRETH is only present when XLEN=32 and reads bits 63-32 of the same instruction counter. The underlying 64-bit counter should never overflow in practice.

Instructions that cause synchronous exceptions, including ECALL and EBREAK, are not considered to retire and hence do not increment the instret CSR.

The following code sequence will read a valid 64-bit cycle counter value into x3:x2, even if the counter overflows its lower half between reading its upper and lower halves.

Listing 1. Sample code for reading the 64-bit cycle counter when XLEN=32.
    again:
        rdcycleh     x3
        rdcycle      x2
        rdcycleh     x4
        bne          x3, x4, again

8.2. "Zihpm" Extension for Hardware Performance Counters

CV32A65X: This extension is not supported.

9. "Zihintntl" Extension for Non-Temporal Locality Hints, Version 1.0

CV32A65X: This extension is not supported.

10. "Zihintpause" Extension for Pause Hint, Version 2.0

CV32A65X: This extension is not supported.

11. "Zimop" Extension for May-Be-Operations, Version 1.0

CV32A65X: This extension is not supported.

11.1. "Zcmop" Compressed May-Be-Operations Extension, Version 1.0

CV32A65X: This extension is not supported.

12. "Zicond" Extension for Integer Conditional Operations, Version 1.0.0

CV32A65X: This extension is not supported.

13. "M" Extension for Integer Multiplication and Division, Version 2.0

This chapter describes the standard integer multiplication and division instruction extension, which is named "M" and contains instructions that multiply or divide values held in two integer registers.

We separate integer multiply and divide out from the base to simplify low-end implementations, or for applications where integer multiply and divide operations are either infrequent or better handled in attached accelerators.

13.1. Multiplication Operations

Diagram

MUL performs an XLEN-bit×XLEN-bit multiplication of rs1 by rs2 and places the lower XLEN bits in the destination register. MULH, MULHU, and MULHSU perform the same multiplication but return the upper XLEN bits of the full 2×XLEN-bit product, for signed×signed, unsigned×unsigned, and rs1×unsigned rs2 multiplication, respectively. If both the high and low bits of the same product are required, then the recommended code sequence is: MULH[[S]U] rdh, rs1, rs2; MUL rdl, rs1, rs2 (source register specifiers must be in same order and rdh cannot be the same as rs1 or rs2). Microarchitectures can then fuse these into a single multiply operation instead of performing two separate multiplies.

MULHSU is used in multi-word signed multiplication to multiply the most-significant word of the multiplicand (which contains the sign bit) with the less-significant words of the multiplier (which are unsigned).

MULW is an RV64 instruction that multiplies the lower 32 bits of the source registers, placing the sign extension of the lower 32 bits of the result into the destination register.

In RV64, MUL can be used to obtain the upper 32 bits of the 64-bit product, but signed arguments must be proper 32-bit signed values, whereas unsigned arguments must have their upper 32 bits clear. If the arguments are not known to be sign- or zero-extended, an alternative is to shift both arguments left by 32 bits, then use MULH[[S]U].

13.2. Division Operations

Diagram

DIV and DIVU perform an XLEN bits by XLEN bits signed and unsigned integer division of rs1 by rs2, rounding towards zero. REM and REMU provide the remainder of the corresponding division operation. For REM, the sign of a nonzero result equals the sign of the dividend.

For both signed and unsigned division, except in the case of overflow, it holds that stem 5737b426f42585386046d1ae40cf6e65.

If both the quotient and remainder are required from the same division, the recommended code sequence is: DIV[U] rdq, rs1, rs2; REM[U] rdr, rs1, rs2 (rdq cannot be the same as rs1 or rs2). Microarchitectures can then fuse these into a single divide operation instead of performing two separate divides.

DIVW and DIVUW are RV64 instructions that divide the lower 32 bits of rs1 by the lower 32 bits of rs2, treating them as signed and unsigned integers respectively, placing the 32-bit quotient in rd, sign-extended to 64 bits. REMW and REMUW are RV64 instructions that provide the corresponding signed and unsigned remainder operations respectively. Both REMW and REMUW always sign-extend the 32-bit result to 64 bits, including on a divide by zero.

The semantics for division by zero and division overflow are summarized in Table 8. The quotient of division by zero has all bits set, and the remainder of division by zero equals the dividend. Signed division overflow occurs only when the most-negative integer is divided by stem 52c1acbf19a35be815dca1fd989d89e4. The quotient of a signed division with overflow is equal to the dividend, and the remainder is zero. Unsigned division overflow cannot occur.

Table 8. Semantics for division by zero and division overflow. L is the width of the operation in bits: XLEN for DIV[U] and REM[U], or 32 for DIV[U]W and REM[U]W.
Condition Dividend Divisor DIVU[W] REMU[W] DIV[W] REM[W]

Division by zero
Overflow (signed only)

stem 7073627e9999e583f5539cb4560a14d7
stem 1f312cf5724cc9327753b6aa6e2b3c6c

0
stem 52c1acbf19a35be815dca1fd989d89e4

stem 2211f7c99ec951c50baf5ffaf8b1b52d
-

stem 7073627e9999e583f5539cb4560a14d7
-

stem 52c1acbf19a35be815dca1fd989d89e4
stem 1f312cf5724cc9327753b6aa6e2b3c6c

stem 7073627e9999e583f5539cb4560a14d7
0

We considered raising exceptions on integer divide by zero, with these exceptions causing a trap in most execution environments. However, this would be the only arithmetic trap in the standard ISA (floating-point exceptions set flags and write default values, but do not cause traps) and would require language implementors to interact with the execution environment’s trap handlers for this case. Further, where language standards mandate that a divide-by-zero exception must cause an immediate control flow change, only a single branch instruction needs to be added to each divide operation, and this branch instruction can be inserted after the divide and should normally be very predictably not taken, adding little runtime overhead.

The value of all bits set is returned for both unsigned and signed divide by zero to simplify the divider circuitry. The value of all 1s is both the natural value to return for unsigned divide, representing the largest unsigned number, and also the natural result for simple unsigned divider implementations. Signed division is often implemented using an unsigned division circuit and specifying the same overflow result simplifies the hardware.

13.3. Zmmul Extension, Version 1.0

The Zmmul extension implements the multiplication subset of the M extension. It adds all of the instructions defined in Section 13.1, namely: MUL, MULH, MULHU, MULHSU, and (for RV64 only) MULW. The encodings are identical to those of the corresponding M-extension instructions. M implies Zmmul.

The Zmmul extension enables low-cost implementations that require multiplication operations but not division. For many microcontroller applications, division operations are too infrequent to justify the cost of divider hardware. By contrast, multiplication operations are more frequent, making the cost of multiplier hardware more justifiable. Simple FPGA soft cores particularly benefit from eliminating division but retaining multiplication, since many FPGAs provide hardwired multipliers but require dividers be implemented in soft logic.

14. "A" Extension for Atomic Instructions, Version 2.1

CV32A65X: This extension is not supported.

15. "Zawrs" Extension for Wait-on-Reservation-Set instructions, Version 1.01

CV32A65X: This extension is not supported.

16. "Zacas" Extension for Atomic Compare-and-Swap (CAS) Instructions, Version 1.0.0

CV32A65X: This extension is not supported.

17. "Zabha" Extension for Byte and Halfword Atomic Memory Operations, Version 1.0

CV32A65X: This extension is not supported.

18. RVWMO Memory Consistency Model, Version 2.0

This chapter defines the RISC-V memory consistency model. A memory consistency model is a set of rules specifying the values that can be returned by loads of memory. RISC-V uses a memory model called "RVWMO" (RISC-V Weak Memory Ordering) which is designed to provide flexibility for architects to build high-performance scalable designs while simultaneously supporting a tractable programming model.

Under RVWMO, code running on a single hart appears to execute in order from the perspective of other memory instructions in the same hart, but memory instructions from another hart may observe the memory instructions from the first hart being executed in a different order. Therefore, multithreaded code may require explicit synchronization to guarantee ordering between memory instructions from different harts. The base RISC-V ISA provides a FENCE instruction for this purpose, described in Section 2.7, while the atomics extension "A" additionally defines load-reserved/store-conditional and atomic read-modify-write instructions.

The standard ISA extension for total store ordering "Ztso" (Chapter 19) augments RVWMO with additional rules specific to those extensions.

The appendices to this specification provide both axiomatic and operational formalizations of the memory consistency model as well as additional explanatory material.

This chapter defines the memory model for regular main memory operations. The interaction of the memory model with I/O memory, instruction fetches, FENCE.I, page table walks, and SFENCE.VMA is not (yet) formalized. Some or all of the above may be formalized in a future revision of this specification. The RV128 base ISA and future ISA extensions such as the V vector and J JIT extensions will need to be incorporated into a future revision as well.

Memory consistency models supporting overlapping memory accesses of different widths simultaneously remain an active area of academic research and are not yet fully understood. The specifics of how memory accesses of different sizes interact under RVWMO are specified to the best of our current abilities, but they are subject to revision should new issues be uncovered.

18.1. Definition of the RVWMO Memory Model

The RVWMO memory model is defined in terms of the global memory order, a total ordering of the memory operations produced by all harts. In general, a multithreaded program has many different possible executions, with each execution having its own corresponding global memory order.

The global memory order is defined over the primitive load and store operations generated by memory instructions. It is then subject to the constraints defined in the rest of this chapter. Any execution satisfying all of the memory model constraints is a legal execution (as far as the memory model is concerned).

18.1.1. Memory Model Primitives

The program order over memory operations reflects the order in which the instructions that generate each load and store are logically laid out in that hart’s dynamic instruction stream; i.e., the order in which a simple in-order processor would execute the instructions of that hart.

Memory-accessing instructions give rise to memory operations. A memory operation can be either a load operation, a store operation, or both simultaneously. All memory operations are single-copy atomic: they can never be observed in a partially complete state.

Among instructions in RV32GC and RV64GC, each aligned memory instruction gives rise to exactly one memory operation, with two exceptions. First, an unsuccessful SC instruction does not give rise to any memory operations. Second, FLD and FSD instructions may each give rise to multiple memory operations if XLEN<64, as stated in [fld_fsd] and clarified below. An aligned AMO gives rise to a single memory operation that is both a load operation and a store operation simultaneously.

Instructions in the RV128 base instruction set and in future ISA extensions such as V (vector) and P (SIMD) may give rise to multiple memory operations. However, the memory model for these extensions has not yet been formalized.

A misaligned load or store instruction may be decomposed into a set of component memory operations of any granularity. An FLD or FSD instruction for which XLEN<64 may also be decomposed into a set of component memory operations of any granularity. The memory operations generated by such instructions are not ordered with respect to each other in program order, but they are ordered normally with respect to the memory operations generated by preceding and subsequent instructions in program order. The atomics extension "A" does not require execution environments to support misaligned atomic instructions at all. However, if misaligned atomics are supported via the misaligned atomicity granule PMA, then AMOs within an atomicity granule are not decomposed, nor are loads and stores defined in the base ISAs, nor are loads and stores of no more than XLEN bits defined in the F, D, and Q extensions.

The decomposition of misaligned memory operations down to byte granularity facilitates emulation on implementations that do not natively support misaligned accesses. Such implementations might, for example, simply iterate over the bytes of a misaligned access one by one.

An LR instruction and an SC instruction are said to be paired if the LR precedes the SC in program order and if there are no other LR or SC instructions in between; the corresponding memory operations are said to be paired as well (except in case of a failed SC, where no store operation is generated). The complete list of conditions determining whether an SC must succeed, may succeed, or must fail is defined in [sec:lrsc].

Load and store operations may also carry one or more ordering annotations from the following set: "acquire-RCpc", "acquire-RCsc", "release-RCpc", and "release-RCsc". An AMO or LR instruction with aq set has an "acquire-RCsc" annotation. An AMO or SC instruction with rl set has a "release-RCsc" annotation. An AMO, LR, or SC instruction with both aq and rl set has both "acquire-RCsc" and "release-RCsc" annotations.

For convenience, we use the term "acquire annotation" to refer to an acquire-RCpc annotation or an acquire-RCsc annotation. Likewise, a "release annotation" refers to a release-RCpc annotation or a release-RCsc annotation. An "RCpc annotation" refers to an acquire-RCpc annotation or a release-RCpc annotation. An RCsc annotation refers to an acquire-RCsc annotation or a release-RCsc annotation.

In the memory model literature, the term "RCpc" stands for release consistency with processor-consistent synchronization operations, and the term "RCsc" stands for release consistency with sequentially consistent synchronization operations.

While there are many different definitions for acquire and release annotations in the literature, in the context of RVWMO these terms are concisely and completely defined by Preserved Program Order rules 5-7.

"RCpc" annotations are currently only used when implicitly assigned to every memory access per the standard extension "Ztso" (Chapter 19). Furthermore, although the ISA does not currently contain native load-acquire or store-release instructions, nor RCpc variants thereof, the RVWMO model itself is designed to be forwards-compatible with the potential addition of any or all of the above into the ISA in a future extension.

18.1.2. Syntactic Dependencies

The definition of the RVWMO memory model depends in part on the notion of a syntactic dependency, defined as follows.

In the context of defining dependencies, a register refers either to an entire general-purpose register, some portion of a CSR, or an entire CSR. The granularity at which dependencies are tracked through CSRs is specific to each CSR and is defined in Section 18.2.

Syntactic dependencies are defined in terms of instructions' source registers, instructions' destination registers, and the way instructions carry a dependency from their source registers to their destination registers. This section provides a general definition of all of these terms; however, Section 18.3 provides a complete listing of the specifics for each instruction.

In general, a register r other than x0 is a source register for an instruction i if any of the following hold:

  • In the opcode of i, rs1, rs2, or rs3 is set to r

  • i is a CSR instruction, and in the opcode of i, csr is set to r, unless i is CSRRW or CSRRWI and rd is set to x0

  • r is a CSR and an implicit source register for i, as defined in Section 18.3

  • r is a CSR that aliases with another source register for i

Memory instructions also further specify which source registers are address source registers and which are data source registers.

In general, a register r other than x0 is a destination register for an instruction i if any of the following hold:

  • In the opcode of i, rd is set to r

  • i is a CSR instruction, and in the opcode of i, csr is set to r, unless i is CSRRS or CSRRC and rs1 is set to x0 or i is CSRRSI or CSRRCI and uimm[4:0] is set to zero.

  • r is a CSR and an implicit destination register for i, as defined in Section 18.3

  • r is a CSR that aliases with another destination register for i

Most non-memory instructions carry a dependency from each of their source registers to each of their destination registers. However, there are exceptions to this rule; see Section 18.3.

Instruction j has a syntactic dependency on instruction i via destination register s of i and source register r of j if either of the following hold:

  • s is the same as r, and no instruction program-ordered between i and j has r as a destination register

  • There is an instruction m program-ordered between i and j such that all of the following hold:

    1. j has a syntactic dependency on m via destination register q and source register r

    2. m has a syntactic dependency on i via destination register s and source register p

    3. m carries a dependency from p to q

Finally, in the definitions that follow, let a and b be two memory operations, and let i and j be the instructions that generate a and b, respectively.

b has a syntactic address dependency on a if r is an address source register for j and j has a syntactic dependency on i via source register r

b has a syntactic data dependency on a if b is a store operation, r is a data source register for j, and j has a syntactic dependency on i via source register r

b has a syntactic control dependency on a if there is an instruction m program-ordered between i and j such that m is a branch or indirect jump and m has a syntactic dependency on i.

Generally speaking, non-AMO load instructions do not have data source registers, and unconditional non-AMO store instructions do not have destination registers. However, a successful SC instruction is considered to have the register specified in rd as a destination register, and hence it is possible for an instruction to have a syntactic dependency on a successful SC instruction that precedes it in program order.

18.1.3. Preserved Program Order

The global memory order for any given execution of a program respects some but not all of each hart’s program order. The subset of program order that must be respected by the global memory order is known as preserved program order.

The complete definition of preserved program order is as follows (and note that AMOs are simultaneously both loads and stores): memory operation a precedes memory operation b in preserved program order (and hence also in the global memory order) if a precedes b in program order, a and b both access regular main memory (rather than I/O regions), and any of the following hold:

  • Overlapping-Address Orderings:

    1. b is a store, and a and b access overlapping memory addresses

    2. a and b are loads, x is a byte read by both a and b, there is no store to x between a and b in program order, and a and b return values for x written by different memory operations

    3. a is generated by an AMO or SC instruction, b is a load, and b returns a value written by a

  • Explicit Synchronization

    1. There is a FENCE instruction that orders a before b

    2. a has an acquire annotation

    3. b has a release annotation

    4. a and b both have RCsc annotations

    5. a is paired with b

  • Syntactic Dependencies

    1. b has a syntactic address dependency on a

    2. b has a syntactic data dependency on a

    3. b is a store, and b has a syntactic control dependency on a

  • Pipeline Dependencies

    1. b is a load, and there exists some store m between a and b in program order such that m has an address or data dependency on a, and b returns a value written by m

    2. b is a store, and there exists some instruction m between a and b in program order such that m has an address dependency on a

18.1.4. Memory Model Axioms

An execution of a RISC-V program obeys the RVWMO memory consistency model only if there exists a global memory order conforming to preserved program order and satisfying the load value axiom, the atomicity axiom, and the progress axiom.

18.1.4.1. Load Value Axiom

Each byte of each load i returns the value written to that byte by the store that is the latest in global memory order among the following stores:

  1. Stores that write that byte and that precede i in the global memory order

  2. Stores that write that byte and that precede i in program order

18.1.4.2. Atomicity Axiom

If r and w are paired load and store operations generated by aligned LR and SC instructions in a hart h, s is a store to byte x, and r returns a value written by s, then s must precede w in the global memory order, and there can be no store from a hart other than h to byte x following s and preceding w in the global memory order.

The Atomicity Axiom theoretically supports LR/SC pairs of different widths and to mismatched addresses, since implementations are permitted to allow SC operations to succeed in such cases. However, in practice, we expect such patterns to be rare, and their use is discouraged.
18.1.4.3. Progress Axiom

No memory operation may be preceded in the global memory order by an infinite sequence of other memory operations.

18.2. CSR Dependency Tracking Granularity

Table 9. Granularities at which syntactic dependencies are tracked through CSRs
Name Portions Tracked as Independent Units Aliases

fflags

Bits 4, 3, 2, 1, 0

fcsr

frm

entire CSR

fcsr

fcsr

Bits 7-5, 4, 3, 2, 1, 0

fflags, frm

Note: read-only CSRs are not listed, as they do not participate in the definition of syntactic dependencies.

18.3. Source and Destination Register Listings

This section provides a concrete listing of the source and destination registers for each instruction. These listings are used in the definition of syntactic dependencies in Section 18.1.2.

The term "accumulating CSR" is used to describe a CSR that is both a source and a destination register, but which carries a dependency only from itself to itself.

Instructions carry a dependency from each source register in the "Source Registers" column to each destination register in the "Destination Registers" column, from each source register in the "Source Registers" column to each CSR in the "Accumulating CSRs" column, and from each CSR in the "Accumulating CSRs" column to itself, except where annotated otherwise.

Key:

  • AAddress source register

  • DData source register

  • † The instruction does not carry a dependency from any source register to any destination register

  • ‡ The instruction carries dependencies from source register(s) to destination register(s) as specified

Table 10. RV32I Base Integer Instruction Set
Source Registers Destination Registers Accumulating CSRs

LUI

rd

AUIPC

rd

JAL

rd

JALR†

rs1

rd

BEQ

rs1, rs2

BNE

rs1, rs2

BLT

rs1, rs2

BGE

rs1, rs2

BLTU

rs1, rs2

BGEU

rs1, rs2

LB †

rs1 A

rd

LH †

rs1 A

rd

LW †

rs1 A

rd

LBU †

rs1 A

rd

LHU †

rs1 A

rd

SB

rs1 A, rs2 D

SH

rs1 A, rs2 D

SW

rs1 A, rs2 D

ADDI

rs1

rd

SLTI

rs1

rd

SLTIU

rs1

rd

XORI

rs1

rd

ORI

rs1

rd

ANDI

rs1

rd

SLLI

rs1

rd

SRLI

rs1

rd

SRAI

rs1

rd

ADD

rs1, rs2

rd

SUB

rs1, rs2

rd

SLL

rs1, rs2

rd

SLT

rs1, rs2

rd

SLTU

rs1, rs2

rd

XOR

rs1, rs2

rd

SRL

rs1, rs2

rd

SRA

rs1, rs2

rd

OR

rs1, rs2

rd

AND

rs1, rs2

rd

FENCE

FENCE.I

ECALL

EBREAK

CSRRW‡

rs1, csr*

rd, csr

*unless rd=x0

CSRRS‡

rs1, csr

rd *, csr

*unless rs1=x0

CSRRC‡

rs1, csr

rd *, csr

*unless rs1=x0

‡ carries a dependency from rs1 to csr and from csr to rd

CSRRWI ‡

csr *

rd, csr

*unless rd=x0

CSRRSI ‡

csr

rd, csr*

*unless uimm[4:0]=0

CSRRCI ‡

csr

rd, csr*

*unless uimm[4:0]=0

‡ carries a dependency from csr to rd
Table 11. RV64I Base Integer Instruction Set
Source Registers Destination Registers Accumulating CSRs

LWU †

rs1 A

rd

LD †

rs1 A

rd

SD

rs1 A, rs2 D

SLLI

rs1

rd

SRLI

rs1

rd

SRAI

rs1

rd

ADDIW

rs1

rd

SLLIW

rs1

rd

SRLIW

rs1

rd

SRAIW

rs1

rd

ADDW

rs1, rs2

rd

SUBW

rs1, rs2

rd

SLLW

rs1, rs2

rd

SRLW

rs1, rs2

rd

SRAW

rs1, rs2

rd
Table 12. RV32M Standard Extension
Source Registers Destination Registers Accumulating CSRs

MUL

rs1, rs2

rd

MULH

rs1, rs2

rd

MULHSU

rs1, rs2

rd

MULHU

rs1, rs2

rd

DIV

rs1, rs2

rd

DIVU

rs1, rs2

rd

REM

rs1, rs2

rd

REMU

rs1, rs2

rd
Table 13. RV64M Standard Extension
Source Registers Destination Registers Accumulating CSRs

MULW

rs1, rs2

rd

DIVW

rs1, rs2

rd

DIVUW

rs1, rs2

rd

REMW

rs1, rs2

rd

REMUW

rs1, rs2

rd
Table 14. RV32A Standard Extension
Source Registers Destination Registers Accumulating CSRs

LR.W†

rs1 A

rd

SC.W†

rs1 A, rs2 D

rd *

* if successful

AMOSWAP.W†

rs1 A, rs2 D

rd

AMOADD.W†

rs1 A, rs2 D

rd

AMOXOR.W†

rs1 A, rs2 D

rd

AMOAND.W†

rs1 A, rs2 D

rd

AMOOR.W†

rs1 A, rs2D

rd

AMOMIN.W†

rs1 A, rs2 D

rd

AMOMAX.W†

rs1 A, rs2 D

rd

AMOMINU.W†

rs1 A, rs2 D

rd

AMOMAXU.W†

rs1 A, rs2 D

rd
Table 15. RV64A Standard Extension
Source Registers Destination Registers Accumulating CSRs

LR.D†

rs1 A

rd

SC.D†

rs1 A, rs2 D

rd *

*if successful

AMOSWAP.D†

rs1 A, rs2 D

rd

AMOADD.D†

rs1 A, rs2 D

rd

AMOXOR.D†

rs1 A, rs2 D

rd

AMOAND.D†

rs1 A, rs2D

rd

AMOOR.D†

rs1 A, rs2D

rd

AMOMIN.D†

rs1 A, rs2D

rd

AMOMAX.D†

rs1 A, rs2D

rd

AMOMINU.D†

rs1 A, rs2D

rd

AMOMAXU.D†

rs1 A, rs2D

rd
Table 16. RV32F Standard Extension
Source Registers Destination Registers Accumulating CSRs

FLW†

rs1 A

rd

FSW

rs1 A, rs2D

FMADD.S

rs1, rs2, rs3, frm*

rd

NV, OF, UF, NX

*if rm=111

FMSUB.S

rs1, rs2, rs3, frm*

rd

NV, OF, UF, NX

*if rm=111

FNMSUB.S

rs1, rs2, rs3, frm*

rd

NV, OF, UF, NX

*if rm=111

FNMADD.S

rs1, rs2, rs3, frm*

rd

NV, OF, UF, NX

*if rm=111

FADD.S

rs1, rs2, frm*

rd

NV, OF, NX

*if rm=111

FSUB.S

rs1, rs2, frm*

rd

NV, OF, NX

*if rm=111

FMUL.S

rs1, rs2, frm*

rd

NV, OF, UF, NX

*if rm=111

FDIV.S

rs1, rs2, frm*

rd

NV, DZ, OF, UF, NX

*if rm=111

FSQRT.S

rs1, frm*

rd

NV, NX

*if rm=111

FSGNJ.S

rs1, rs2

rd

FSGNJN.S

rs1, rs2

rd

FSGNJX.S

rs1, rs2

rd

FMIN.S

rs1, rs2

rd

NV

FMAX.S

rs1, rs2

rd

NV

FCVT.W.S

rs1, frm*

rd

NV, NX

*if rm=111

FCVT.WU.S

rs1, frm*

rd

NV, NX

*if rm=111

FMV.X.W

rs1

rd

FEQ.S

rs1, rs2

rd

NV

FLT.S

rs1, rs2

rd

NV

FLE.S

rs1, rs2

rd

NV

FCLASS.S

rs1

rd

FCVT.S.W

rs1, frm*

rd

NX

*if rm=111

FCVT.S.WU

rs1, frm*

rd

NX

*if rm=111

FMV.W.X

rs1

rd
Table 17. RV64F Standard Extension
Source Registers Destination Registers Accumulating CSRs

FCVT.L.S

rs1, frm*

rd

NV, NX

*if rm=111

FCVT.LU.S

rs1, frm*

rd

NV, NX

*if rm=111

FCVT.S.L

rs1, frm*

rd

NX

*if rm=111

FCVT.S.LU

rs1, frm*

rd

NX

*if rm=111
Table 18. RV32D Standard Extension
Source Registers Destination Registers Accumulating CSRs

FLD†

rs1 A

rd

FSD

rs1 A, rs2D

FMADD.D

rs1, rs2, rs3, frm*

rd

NV, OF, UF, NX

*if rm=111

FMSUB.D

rs1, rs2, rs3, frm*

rd

NV, OF, UF, NX

*if rm=111

FNMSUB.D

rs1, rs2, rs3, frm*

rd

NV, OF, UF, NX

*if rm=111

FNMADD.D

rs1, rs2, rs3, frm*

rd

NV, OF, UF, NX

*if rm=111

FADD.D

rs1, rs2, frm*

rd

NV, OF, NX

*if rm=111

FSUB.D

rs1, rs2, frm*

rd

NV, OF, NX

*if rm=111

FMUL.D

rs1, rs2, frm*

rd

NV, OF, UF, NX

*if rm=111

FDIV.D

rs1, rs2, frm*

rd

NV, DZ, OF, UF, NX

*if rm=111

FSQRT.D

rs1, frm*

rd

NV, NX

*if rm=111

FSGNJ.D

rs1, rs2

rd

FSGNJN.D

rs1, rs2

rd

FSGNJX.D

rs1, rs2

rd

FMIN.D

rs1, rs2

rd

NV

FMAX.D

rs1, rs2

rd

NV

FCVT.S.D

rs1, frm*

rd

NV, OF, UF, NX

*if rm=111

FCVT.D.S

rs1

rd

NV

FEQ.D

rs1, rs2

rd

NV

FLT.D

rs1, rs2

rd

NV

FLE.D

rs1, rs2

rd

NV

FCLASS.D

rs1

rd

FCVT.W.D

rs1,*

rd

NV, NX

*if rm=111

FCVT.WU.D

rs1, frm*

rd

NV, NX

*if rm=111

FCVT.D.W

rs1

rd

FCVT.D.WU

rs1

rd
Table 19. RV64D Standard Extension
Source Registers Destination Registers Accumulating CSRs

FCVT.L.D

rs1, frm*

rd

NV, NX

*if rm=111

FCVT.LU.D

rs1, frm*

rd

NV, NX

*if rm=111

FMV.X.D

rs1

rd

FCVT.D.L

rs1, frm*

rd

NX

*if rm=111

FCVT.D.LU

rs1, frm*

rd

NX

*if rm=111

FMV.D.X

rs1

rd

19. "Ztso" Extension for Total Store Ordering, Version 1.0

CV32A65X: This extension is not supported.

20. "CMO" Extensions for Base Cache Management Operation ISA, Version 1.0.0

CV32A65X: These extensions are not supported.

21. "F" Extension for Single-Precision Floating-Point, Version 2.2

CV32A65X: This extension is not supported.

22. "D" Extension for Double-Precision Floating-Point, Version 2.2

CV32A65X: This extension is not supported.

23. "Q" Extension for Quad-Precision Floating-Point, Version 2.2

CV32A65X: This extension is not supported.

24. "Zfh" and "Zfhmin" Extensions for Half-Precision Floating-Point, Version 1.0

CV32A65X: These extensions are not supported.

25. "BF16" Extensions for for BFloat16-precision Floating-Point, Version 1.0

CV32A65X: These extensions are not supported.

26. "Zfa" Extension for Additional Floating-Point Instructions, Version 1.0

CV32A65X: This extension is not supported.

27. "Zfinx", "Zdinx", "Zhinx", "Zhinxmin" Extensions for Floating-Point in Integer Registers, Version 1.0

CV32A65X: These extensions are not supported.

28. "C" Extension for Compressed Instructions, Version 2.0

This chapter describes the RISC-V standard compressed instruction-set extension, named "C", which reduces static and dynamic code size by adding short 16-bit instruction encodings for common operations. The C extension can be added to any of the base ISAs (RV32, RV64, RV128), and we use the generic term "RVC" to cover any of these. Typically, 50%-60% of the RISC-V instructions in a program can be replaced with RVC instructions, resulting in a 25%-30% code-size reduction.

28.1. Overview

RVC uses a simple compression scheme that offers shorter 16-bit versions of common 32-bit RISC-V instructions when:

  • the immediate or address offset is small, or

  • one of the registers is the zero register (x0), the ABI link register (x1), or the ABI stack pointer (x2), or

  • the destination register and the first source register are identical, or

  • the registers used are the 8 most popular ones.

The C extension is compatible with all other standard instruction extensions. The C extension allows 16-bit instructions to be freely intermixed with 32-bit instructions, with the latter now able to start on any 16-bit boundary, i.e., IALIGN=16. With the addition of the C extension, no instructions can raise instruction-address-misaligned exceptions.

Removing the 32-bit alignment constraint on the original 32-bit instructions allows significantly greater code density.

The compressed instruction encodings are mostly common across RV32C, RV64C, and RV128C, but as shown in Table 34, a few opcodes are used for different purposes depending on base ISA. For example, the wider address-space RV64C and RV128C variants require additional opcodes to compress loads and stores of 64-bit integer values, while RV32C uses the same opcodes to compress loads and stores of single-precision floating-point values. Similarly, RV128C requires additional opcodes to capture loads and stores of 128-bit integer values, while these same opcodes are used for loads and stores of double-precision floating-point values in RV32C and RV64C. If the C extension is implemented, the appropriate compressed floating-point load and store instructions must be provided whenever the relevant standard floating-point extension (F and/or D) is also implemented. In addition, RV32C includes a compressed jump and link instruction to compress short-range subroutine calls, where the same opcode is used to compress ADDIW for RV64C and RV128C.

Double-precision loads and stores are a significant fraction of static and dynamic instructions, hence the motivation to include them in the RV32C and RV64C encoding.

Although single-precision loads and stores are not a significant source of static or dynamic compression for benchmarks compiled for the currently supported ABIs, for microcontrollers that only provide hardware single-precision floating-point units and have an ABI that only supports single-precision floating-point numbers, the single-precision loads and stores will be used at least as frequently as double-precision loads and stores in the measured benchmarks. Hence, the motivation to provide compressed support for these in RV32C.

Short-range subroutine calls are more likely in small binaries for microcontrollers, hence the motivation to include these in RV32C.

Although reusing opcodes for different purposes for different base ISAs adds some complexity to documentation, the impact on implementation complexity is small even for designs that support multiple base ISAs. The compressed floating-point load and store variants use the same instruction format with the same register specifiers as the wider integer loads and stores.

RVC was designed under the constraint that each RVC instruction expands into a single 32-bit instruction in either the base ISA (RV32I/E, RV64I/E, or RV128I) or the F and D standard extensions where present. Adopting this constraint has two main benefits:

  • Hardware designs can simply expand RVC instructions during decode, simplifying verification and minimizing modifications to existing microarchitectures.

  • Compilers can be unaware of the RVC extension and leave code compression to the assembler and linker, although a compression-aware compiler will generally be able to produce better results.

We felt the multiple complexity reductions of a simple one-one mapping between C and base IFD instructions far outweighed the potential gains of a slightly denser encoding that added additional instructions only supported in the C extension, or that allowed encoding of multiple IFD instructions in one C instruction.

It is important to note that the C extension is not designed to be a stand-alone ISA, and is meant to be used alongside a base ISA.

Variable-length instruction sets have long been used to improve code density. For example, the IBM Stretch (Buchholz, 1962), developed in the late 1950s, had an ISA with 32-bit and 64-bit instructions, where some of the 32-bit instructions were compressed versions of the full 64-bit instructions. Stretch also employed the concept of limiting the set of registers that were addressable in some of the shorter instruction formats, with short branch instructions that could only refer to one of the index registers. The later IBM 360 architecture (Amdahl et al., 1964) supported a simple variable-length instruction encoding with 16-bit, 32-bit, or 48-bit instruction formats.

In 1963, CDC introduced the Cray-designed CDC 6600 (Thornton, 1965), a precursor to RISC architectures, that introduced a register-rich load-store architecture with instructions of two lengths, 15-bits and 30-bits. The later Cray-1 design used a very similar instruction format, with 16-bit and 32-bit instruction lengths.

The initial RISC ISAs from the 1980s all picked performance over code size, which was reasonable for a workstation environment, but not for embedded systems. Hence, both ARM and MIPS subsequently made versions of the ISAs that offered smaller code size by offering an alternative 16-bit wide instruction set instead of the standard 32-bit wide instructions. The compressed RISC ISAs reduced code size relative to their starting points by about 25-30%, yielding code that was significantly smaller than 80x86. This result surprised some, as their intuition was that the variable-length CISC ISA should be smaller than RISC ISAs that offered only 16-bit and 32-bit formats.

Since the original RISC ISAs did not leave sufficient opcode space free to include these unplanned compressed instructions, they were instead developed as complete new ISAs. This meant compilers needed different code generators for the separate compressed ISAs. The first compressed RISC ISA extensions (e.g., ARM Thumb and MIPS16) used only a fixed 16-bit instruction size, which gave good reductions in static code size but caused an increase in dynamic instruction count, which led to lower performance compared to the original fixed-width 32-bit instruction size. This led to the development of a second generation of compressed RISC ISA designs with mixed 16-bit and 32-bit instruction lengths (e.g., ARM Thumb2, microMIPS, PowerPC VLE), so that performance was similar to pure 32-bit instructions but with significant code size savings. Unfortunately, these different generations of compressed ISAs are incompatible with each other and with the original uncompressed ISA, leading to significant complexity in documentation, implementations, and software tools support.

Of the commonly used 64-bit ISAs, only PowerPC and microMIPS currently supports a compressed instruction format. It is surprising that the most popular 64-bit ISA for mobile platforms (ARM v8) does not include a compressed instruction format given that static code size and dynamic instruction fetch bandwidth are important metrics. Although static code size is not a major concern in larger systems, instruction fetch bandwidth can be a major bottleneck in servers running commercial workloads, which often have a large instruction working set.

Benefiting from 25 years of hindsight, RISC-V was designed to support compressed instructions from the outset, leaving enough opcode space for RVC to be added as a simple extension on top of the base ISA (along with many other extensions). The philosophy of RVC is to reduce code size for embedded applications and to improve performance and energy-efficiency for all applications due to fewer misses in the instruction cache. Waterman shows that RVC fetches 25%-30% fewer instruction bits, which reduces instruction cache misses by 20%-25%, or roughly the same performance impact as doubling the instruction cache size. (Waterman, 2011)

28.2. Compressed Instruction Formats

Table 20 shows the nine compressed instruction formats. CR, CI, and CSS can use any of the 32 RVI registers, but CIW, CL, CS, CA, and CB are limited to just 8 of them. Table 21 lists these popular registers, which correspond to registers x8 to x15. Note that there is a separate version of load and store instructions that use the stack pointer as the base address register, since saving to and restoring from the stack are so prevalent, and that they use the CI and CSS formats to allow access to all 32 data registers. CIW supplies an 8-bit immediate for the ADDI4SPN instruction.

The RISC-V ABI was changed to make the frequently used registers map to registers 'x8-x15'. This simplifies the decompression decoder by having a contiguous naturally aligned set of register numbers, and is also compatible with the RV32E and RV64E base ISAs, which only have 16 integer registers.

Compressed register-based floating-point loads and stores also use the CL and CS formats respectively, with the eight registers mapping to f8 to f15.

The standard RISC-V calling convention maps the most frequently used floating-point registers to registers f8 to f15, which allows the same register decompression decoding as for integer register numbers.

The formats were designed to keep bits for the two register source specifiers in the same place in all instructions, while the destination register field can move. When the full 5-bit destination register specifier is present, it is in the same place as in the 32-bit RISC-V encoding. Where immediates are sign-extended, the sign extension is always from bit 12. Immediate fields have been scrambled, as in the base specification, to reduce the number of immediate muxes required.

The immediate fields are scrambled in the instruction formats instead of in sequential order so that as many bits as possible are in the same position in every instruction, thereby simplifying implementations.

For many RVC instructions, zero-valued immediates are disallowed and x0 is not a valid 5-bit register specifier. These restrictions free up encoding space for other instructions requiring fewer operand bits.

Table 20. Compressed 16-bit RVC instruction formats

Format

Meaning

CR

Register

CI

Immediate

CSS

Stack-relative Store

CIW

Wide Immediate

CL

Load

CS

Store

CA

Arithmetic

CB

Branch/Arithmetic

CJ

Jump

15 14 13 12

11 10 9 8 7

6 5 4 3 2

1 0

funct4

rd/rs1

rs2

op

funct3

imm

rd/rs1

imm

op

funct3

imm

rs2

op

funct3

imm

rd′

op

funct3

imm

rs1′

imm

rd′

op

funct3

imm

rs1′

imm

rs2′

op

funct6

rd′/rs1′

funct2

rs2′

op

funct3

offset

rd′/rs1′

offset

op

funct3

jump target

op
Table 21. Registers specified by the three-bit rs1′, rs2′, and rd′ fields of the CIW, CL, CS, CA, and CB formats.

RVC Register Number

Integer Register Number

Integer Register ABI Name

Floating-Point Register Number

Floating-Point Register ABI Name

000

001

010

011

100

101

110

111

x8

x9

x10

x11

x12

x13

x14

x15

s0

s1

a0

a1

a2

a3

a4

a5

f8

f9

f10

f11

f12

f13

f14

f15

fs0

fs1

fa0

fa1

fa2

fa3

fa4

fa5

28.3. Load and Store Instructions

To increase the reach of 16-bit instructions, data-transfer instructions use zero-extended immediates that are scaled by the size of the data in bytes: ×4 for words, ×8 for double words, and ×16 for quad words.

RVC provides two variants of loads and stores. One uses the ABI stack pointer, x2, as the base address and can target any data register. The other can reference one of 8 base address registers and one of 8 data registers.

28.3.1. Stack-Pointer-Based Loads and Stores

Diagram

These instructions use the CI format.

C.LWSP loads a 32-bit value from memory into register rd. It computes an effective address by adding the zero-extended offset, scaled by 4, to the stack pointer, x2. It expands to lw rd, offset(x2). C.LWSP is only valid when rd≠x0 the code points with rd=x0 are reserved.

C.LDSP is an RV64C/RV128C-only instruction that loads a 64-bit value from memory into register rd. It computes its effective address by adding the zero-extended offset, scaled by 8, to the stack pointer, x2. It expands to ld rd, offset(x2). C.LDSP is only valid when rd≠x0 the code points with rd=x0 are reserved.

C.LQSP is an RV128C-only instruction that loads a 128-bit value from memory into register rd. It computes its effective address by adding the zero-extended offset, scaled by 16, to the stack pointer, x2. It expands to lq rd, offset(x2). C.LQSP is only valid when rd≠x0 the code points with rd=x0 are reserved.

C.FLWSP is an RV32FC-only instruction that loads a single-precision floating-point value from memory into floating-point register rd. It computes its effective address by adding the zero-extended offset, scaled by 4, to the stack pointer, x2. It expands to flw rd, offset(x2).

C.FLDSP is an RV32DC/RV64DC-only instruction that loads a double-precision floating-point value from memory into floating-point register rd. It computes its effective address by adding the zero-extended offset, scaled by 8, to the stack pointer, x2. It expands to fld rd, offset(x2).

Diagram

These instructions use the CSS format.

C.SWSP stores a 32-bit value in register rs2 to memory. It computes an effective address by adding the zero-extended offset, scaled by 4, to the stack pointer, x2. It expands to sw rs2, offset(x2).

C.SDSP is an RV64C/RV128C-only instruction that stores a 64-bit value in register rs2 to memory. It computes an effective address by adding the zero-extended offset, scaled by 8, to the stack pointer, x2. It expands to sd rs2, offset(x2).

C.SQSP is an RV128C-only instruction that stores a 128-bit value in register rs2 to memory. It computes an effective address by adding the zero-extended offset, scaled by 16, to the stack pointer, x2. It expands to sq rs2, offset(x2).

C.FSWSP is an RV32FC-only instruction that stores a single-precision floating-point value in floating-point register rs2 to memory. It computes an effective address by adding the zero-extended offset, scaled by 4, to the stack pointer, x2. It expands to fsw rs2, offset(x2).

C.FSDSP is an RV32DC/RV64DC-only instruction that stores a double-precision floating-point value in floating-point register rs2 to memory. It computes an effective address by adding the zero-extended offset, scaled by 8, to the stack pointer, x2. It expands to fsd rs2, offset(x2).

Register save/restore code at function entry/exit represents a significant portion of static code size. The stack-pointer-based compressed loads and stores in RVC are effective at reducing the save/restore static code size by a factor of 2 while improving performance by reducing dynamic instruction bandwidth.

A common mechanism used in other ISAs to further reduce save/restore code size is load-multiple and store-multiple instructions. We considered adopting these for RISC-V but noted the following drawbacks to these instructions:

  • These instructions complicate processor implementations.

  • For virtual memory systems, some data accesses could be resident in physical memory and some could not, which requires a new restart mechanism for partially executed instructions.

  • Unlike the rest of the RVC instructions, there is no IFD equivalent to Load Multiple and Store Multiple.

  • Unlike the rest of the RVC instructions, the compiler would have to be aware of these instructions to both generate the instructions and to allocate registers in an order to maximize the chances of the them being saved and stored, since they would be saved and restored in sequential order.

  • Simple microarchitectural implementations will constrain how other instructions can be scheduled around the load and store multiple instructions, leading to a potential performance loss.

  • The desire for sequential register allocation might conflict with the featured registers selected for the CIW, CL, CS, CA, and CB formats.

Furthermore, much of the gains can be realized in software by replacing prologue and epilogue code with subroutine calls to common prologue and epilogue code, a technique described in Section 5.6 of (Waterman, 2016).

While reasonable architects might come to different conclusions, we decided to omit load and store multiple and instead use the software-only approach of calling save/restore millicode routines to attain the greatest code size reduction.

28.3.2. Register-Based Loads and Stores

Diagram

These instructions use the CL format.

C.LW loads a 32-bit value from memory into register rd′. It computes an effective address by adding the zero-extended offset, scaled by 4, to the base address in register rs1′. It expands to lw rd′, offset(rs1′).

C.LD is an RV64C/RV128C-only instruction that loads a 64-bit value from memory into register rd′. It computes an effective address by adding the zero-extended offset, scaled by 8, to the base address in register rs1′. It expands to ld rd′, offset(rs1′).

C.LQ is an RV128C-only instruction that loads a 128-bit value from memory into register rd′. It computes an effective address by adding the zero-extended offset, scaled by 16, to the base address in register rs1′. It expands to lq rd′, offset(rs1′).

C.FLW is an RV32FC-only instruction that loads a single-precision floating-point value from memory into floating-point register rd′. It computes an effective address by adding the zero-extended offset, scaled by 4, to the base address in register rs1′. It expands to flw rd′, offset(rs1′).

C.FLD is an RV32DC/RV64DC-only instruction that loads a double-precision floating-point value from memory into floating-point register rd′. It computes an effective address by adding the zero-extended offset, scaled by 8, to the base address in register rs1′. It expands to fld rd′, offset(rs1′).

Diagram

These instructions use the CS format.

C.SW stores a 32-bit value in register rs2′ to memory. It computes an effective address by adding the zero-extended offset, scaled by 4, to the base address in register rs1′. It expands to sw rs2′, offset(rs1′).

C.SD is an RV64C/RV128C-only instruction that stores a 64-bit value in register rs2′ to memory. It computes an effective address by adding the zero-extended offset, scaled by 8, to the base address in register rs1′. It expands to sd rs2′, offset(rs1′).

C.SQ is an RV128C-only instruction that stores a 128-bit value in register rs2′ to memory. It computes an effective address by adding the zero-extended offset, scaled by 16, to the base address in register rs1′. It expands to sq rs2′, offset(rs1′).

C.FSW is an RV32FC-only instruction that stores a single-precision floating-point value in floating-point register rs2′ to memory. It computes an effective address by adding the zero-extended offset, scaled by 4, to the base address in register rs1′. It expands to fsw rs2′, offset(rs1′).

C.FSD is an RV32DC/RV64DC-only instruction that stores a double-precision floating-point value in floating-point register rs2′ to memory. It computes an effective address by adding the zero-extended offset, scaled by 8, to the base address in register rs1′. It expands to fsd rs2′, offset(rs1′).

28.4. Control Transfer Instructions

RVC provides unconditional jump instructions and conditional branch instructions. As with base RVI instructions, the offsets of all RVC control transfer instructions are in multiples of 2 bytes.

Diagram

These instructions use the CJ format.

C.J performs an unconditional control transfer. The offset is sign-extended and added to the pc to form the jump target address. C.J can therefore target a ±2 KiB range. C.J expands to jal x0, offset.

C.JAL is an RV32C-only instruction that performs the same operation as C.J, but additionally writes the address of the instruction following the jump (pc+2) to the link register, x1. C.JAL expands to jal x1, offset.

Diagram

These instructions use the CR format.

C.JR (jump register) performs an unconditional control transfer to the address in register rs1. C.JR expands to jalr x0, 0(rs1). C.JR is only valid when stem 7911d76fb533bd8710d50a065c53f749; the code point with stem 89ec63a1940db41477e8738c3a1f7b05 is reserved.

C.JALR (jump and link register) performs the same operation as C.JR, but additionally writes the address of the instruction following the jump (pc+2) to the link register, x1. C.JALR expands to jalr x1, 0(rs1). C.JALR is only valid when stem 7911d76fb533bd8710d50a065c53f749; the code point with stem 89ec63a1940db41477e8738c3a1f7b05 corresponds to the C.EBREAK instruction.

Strictly speaking, C.JALR does not expand exactly to a base RVI instruction as the value added to the PC to form the link address is 2 rather than 4 as in the base ISA, but supporting both offsets of 2 and 4 bytes is only a very minor change to the base microarchitecture.
Diagram

These instructions use the CB format.

C.BEQZ performs conditional control transfers. The offset is sign-extended and added to the pc to form the branch target address. It can therefore target a ±256 B range. C.BEQZ takes the branch if the value in register rs1′ is zero. It expands to beq rs1′, x0, offset.

C.BNEZ is defined analogously, but it takes the branch if rs1′ contains a nonzero value. It expands to bne rs1′, x0, offset.

28.5. Integer Computational Instructions

RVC provides several instructions for integer arithmetic and constant generation.

28.5.1. Integer Constant-Generation Instructions

The two constant-generation instructions both use the CI instruction format and can target any integer register.

Diagram

C.LI loads the sign-extended 6-bit immediate, imm, into register rd. C.LI expands into addi rd, x0, imm. C.LI is only valid when rd≠x0; the code points with rd=x0 encode HINTs.

C.LUI loads the non-zero 6-bit immediate field into bits 17–12 of the destination register, clears the bottom 12 bits, and sign-extends bit 17 into all higher bits of the destination. C.LUI expands into lui rd, imm. C.LUI is only valid when stem 966f16482ee6eca6c3e071e1950d4dd4, and when the immediate is not equal to zero. The code points with imm=0 are reserved; the remaining code points with rd=x0 are HINTs; and the remaining code points with rd=x2 correspond to the C.ADDI16SP instruction.

28.5.2. Integer Register-Immediate Operations

These integer register-immediate operations are encoded in the CI format and perform operations on an integer register and a 6-bit immediate.

Diagram

C.ADDI adds the non-zero sign-extended 6-bit immediate to the value in register rd then writes the result to rd. C.ADDI expands into addi rd, rd, imm. C.ADDI is only valid when rd≠x0 and imm≠0. The code points with rd=x0 encode the C.NOP instruction; the remaining code points with imm=0 encode HINTs.

C.ADDIW is an RV64C/RV128C-only instruction that performs the same computation but produces a 32-bit result, then sign-extends result to 64 bits. C.ADDIW expands into addiw rd, rd, imm. The immediate can be zero for C.ADDIW, where this corresponds to sext.w rd. C.ADDIW is only valid when rd≠x0; the code points with rd=x0 are reserved.

C.ADDI16SP (add immediate to stack pointer) shares the opcode with C.LUI, but has a destination field of x2. C.ADDI16SP adds the non-zero sign-extended 6-bit immediate to the value in the stack pointer (sp=x2), where the immediate is scaled to represent multiples of 16 in the range [-512, 496]. C.ADDI16SP is used to adjust the stack pointer in procedure prologues and epilogues. It expands into addi x2, x2, nzimm[9:4]. C.ADDI16SP is only valid when nzimm≠0; the code point with nzimm=0 is reserved.

In the standard RISC-V calling convention, the stack pointer sp is always 16-byte aligned.
Diagram

C.ADDI4SPN (add immediate to stack pointer, non-destructive) is a CIW-format instruction that adds a zero-extended non-zero immediate, scaled by 4, to the stack pointer, x2, and writes the result to rd′. This instruction is used to generate pointers to stack-allocated variables, and expands to addi rd′, x2, nzuimm[9:2]. C.ADDI4SPN is only valid when nzuimm≠0; the code points with nzuimm=0 are reserved.

Diagram

C.SLLI is a CI-format instruction that performs a logical left shift of the value in register rd then writes the result to rd. The shift amount is encoded in the shamt field. For RV128C, a shift amount of zero is used to encode a shift of 64. C.SLLI expands into slli rd, rd, shamt[5:0], except for RV128C with shamt=0, which expands to slli rd, rd, 64.

For RV32C, shamt[5] must be zero; the code points with shamt[5]=1 are designated for custom extensions. For RV32C and RV64C, the shift amount must be non-zero; the code points with shamt=0 are HINTs. For all base ISAs, the code points with rd=x0 are HINTs, except those with shamt[5]=1 in RV32C.

Diagram

C.SRLI is a CB-format instruction that performs a logical right shift of the value in register rd′ then writes the result to rd′. The shift amount is encoded in the shamt field. For RV128C, a shift amount of zero is used to encode a shift of 64. Furthermore, the shift amount is sign-extended for RV128C, and so the legal shift amounts are 1-31, 64, and 96-127. C.SRLI expands into srli rd′, rd′, shamt, except for RV128C with shamt=0, which expands to srli rd′, rd′, 64.

For RV32C, shamt[5] must be zero; the code points with shamt[5]=1 are designated for custom extensions. For RV32C and RV64C, the shift amount must be non-zero; the code points with shamt=0 are HINTs.

C.SRAI is defined analogously to C.SRLI, but instead performs an arithmetic right shift. C.SRAI expands to srai rd′, rd′, shamt.

Left shifts are usually more frequent than right shifts, as left shifts are frequently used to scale address values. Right shifts have therefore been granted less encoding space and are placed in an encoding quadrant where all other immediates are sign-extended. For RV128, the decision was made to have the 6-bit shift-amount immediate also be sign-extended. Apart from reducing the decode complexity, we believe right-shift amounts of 96-127 will be more useful than 64-95, to allow extraction of tags located in the high portions of 128-bit address pointers. We note that RV128C will not be frozen at the same point as RV32C and RV64C, to allow evaluation of typical usage of 128-bit address-space codes.
Diagram

C.ANDI is a CB-format instruction that computes the bitwise AND of the value in register rd′ and the sign-extended 6-bit immediate, then writes the result to rd′. C.ANDI expands to andi rd′, rd′, imm.

28.5.3. Integer Register-Register Operations

Diagram

These instructions use the CR format.

C.MV copies the value in register rs2 into register rd. C.MV expands into add rd, x0, rs2. C.MV is only valid when rs2≠x0 the code points with rs2=x0 correspond to the C.JR instruction. The code points with rs2≠x0 and rd=x0 are HINTs.

C.MV expands to a different instruction than the canonical MV pseudoinstruction, which instead uses ADDI. Implementations that handle MV specially, e.g. using register-renaming hardware, may find it more convenient to expand C.MV to MV instead of ADD, at slight additional hardware cost.

C.ADD adds the values in registers rd and rs2 and writes the result to register rd. C.ADD expands into add rd, rd, rs2. C.ADD is only valid when rs2≠x0 the code points with rs2=x0 correspond to the C.JALR and C.EBREAK instructions. The code points with rs2≠x0 and rd=x0 are HINTs.

Diagram

These instructions use the CA format.

C.AND computes the bitwise AND of the values in registers rd′ and rs2′, then writes the result to register rd′. C.AND expands into and rd′, rd′, rs2′.

C.OR computes the bitwise OR of the values in registers rd′ and rs2′, then writes the result to register rd′. C.OR expands into or rd′, rd′, rs2′.

C.XOR computes the bitwise XOR of the values in registers rd′ and rs2′, then writes the result to register rd′. C.XOR expands into xor rd′, rd′, rs2′.

C.SUB subtracts the value in register rs2′ from the value in register rd′, then writes the result to register rd′. C.SUB expands into sub rd′, rd′, rs2′.

C.ADDW is an RV64C/RV128C-only instruction that adds the values in registers rd′ and rs2′, then sign-extends the lower 32 bits of the sum before writing the result to register rd′. C.ADDW expands into addw rd′, rd′, rs2′.

C.SUBW is an RV64C/RV128C-only instruction that subtracts the value in register rs2′ from the value in register rd′, then sign-extends the lower 32 bits of the difference before writing the result to register rd′. C.SUBW expands into subw rd′, rd′, rs2′.

This group of six instructions do not provide large savings individually, but do not occupy much encoding space and are straightforward to implement, and as a group provide a worthwhile improvement in static and dynamic compression.

28.5.4. Defined Illegal Instruction

Diagram

A 16-bit instruction with all bits zero is permanently reserved as an illegal instruction.

We reserve all-zero instructions to be illegal instructions to help trap attempts to execute zero-ed or non-existent portions of the memory space. The all-zero value should not be redefined in any non-standard extension. Similarly, we reserve instructions with all bits set to 1 (corresponding to very long instructions in the RISC-V variable-length encoding scheme) as illegal to capture another common value seen in non-existent memory regions.

28.5.5. NOP Instruction

Diagram

C.NOP is a CI-format instruction that does not change any user-visible state, except for advancing the pc and incrementing any applicable performance counters. C.NOP expands to nop. C.NOP is only valid when imm=0; the code points with imm≠0 encode HINTs.

28.5.6. Breakpoint Instruction

Diagram

Debuggers can use the C.EBREAK instruction, which expands to ebreak, to cause control to be transferred back to the debugging environment. C.EBREAK shares the opcode with the C.ADD instruction, but with rd and rs2 both zero, thus can also use the CR format.

28.6. Usage of C Instructions in LR/SC Sequences

On implementations that support the C extension, compressed forms of the I instructions permitted inside constrained LR/SC sequences, as described in [sec:lrscseq], are also permitted inside constrained LR/SC sequences.

The implication is that any implementation that claims to support both the A and C extensions must ensure that LR/SC sequences containing valid C instructions will eventually complete.

28.7. HINT Instructions

A portion of the RVC encoding space is reserved for microarchitectural HINTs. Like the HINTs in the RV32I base ISA (see HINT Instructions), these instructions do not modify any architectural state, except for advancing the pc and any applicable performance counters. HINTs are executed as no-ops on implementations that ignore them.

RVC HINTs are encoded as computational instructions that do not modify the architectural state, either because rd=x0 (e.g. C.ADD x0, t0), or because rd is overwritten with a copy of itself (e.g. C.ADDI t0, 0).

This HINT encoding has been chosen so that simple implementations can ignore HINTs altogether, and instead execute a HINT as a regular computational instruction that happens not to mutate the architectural state.

RVC HINTs do not necessarily expand to their RVI HINT counterparts. For example, C.ADD x0, a0 might not encode the same HINT as ADD x0, x0, a0.

The primary reason to not require an RVC HINT to expand to an RVI HINT is that HINTs are unlikely to be compressible in the same manner as the underlying computational instruction. Also, decoupling the RVC and RVI HINT mappings allows the scarce RVC HINT space to be allocated to the most popular HINTs, and in particular, to HINTs that are amenable to macro-op fusion.

Table 32 lists all RVC HINT code points. For RV32C, 78% of the HINT space is reserved for standard HINTs. The remainder of the HINT space is designated for custom HINTs; no standard HINTs will ever be defined in this subspace.

Table 22. RVC HINT instructions.
Instruction Constraints Code Points Purpose

C.NOP

imm≠0

63

Designated for future standard use

C.ADDI

rd≠x0, imm=0

31

C.LI

rd=x0

64

C.LUI

rd=x0, imm≠0

63

C.MV

rd=x0, rs2≠x0

31

C.ADD

rd=x0, rs2≠x0, rs2≠x2-x5

27

C.ADD

rd=x0, rs2≠x2-x5

4

(rs2=x2) C.NTL.P1 (rs2=x3) C.NTL.PALL (rs2=x4) C.NTL.S1 (rs2=x5) C.NTL.ALL

C.SLLI

rd=x0, imm≠0

31 (RV32), 63 (RV64/128)

Designated for custom use

C.SLLI64

rd=x0

1

C.SLLI64

rd≠x0, RV32 and RV64 only

31

C.SRLI64

RV32 and RV64 only

8

C.SRAI64

RV32 and RV64 only

8

28.8. RVC Instruction Set Listings

Table 23 shows a map of the major opcodes for RVC. Each row of the table corresponds to one quadrant of the encoding space. The last quadrant, which has the two least-significant bits set, corresponds to instructions wider than 16 bits, including those in the base ISAs. Several instructions are only valid for certain operands; when invalid, they are marked either RES to indicate that the opcode is reserved for future standard extensions; Custom to indicate that the opcode is designated for custom extensions; or HINT to indicate that the opcode is reserved for microarchitectural hints (see Section 18.7).

Table 23. RVC opcode map instructions.

inst[15:13]
inst[1:0]

000

001

010

011

100

101

110

111

00

ADDI4SPN

FLD
FLD
LQ

LW

FLW
LD
LD

Reserved

FSD
FSD
SQ

SW

FSW
SD
SD

RV32
RV64
RV128

01

ADDI

JAL
ADDIW
ADDIW

LI

LUI/ADDI16SP

MISC-ALU

J

BEQZ

BNEZ

RV32
RV64
RV128

10

SLLI

FLDSP
FLDSP
LQSP

LWSP

FLWSP
LDSP
LDSP

J[AL]R/MV/ADD

FSDSP
FSDSP
SQSP

SWSP

FSWSP
SDSP
SDSP

RV32
RV64
RV128

11

>16b

Figure 2, Figure 3, and Figure 4 list the RVC instructions.

Diagram
Figure 2. Instruction listing for RVC, Quadrant 0
Diagram
Figure 3. Instruction listing for RVC, Quadrant 1
Diagram
Figure 4. Instruction listing for RVC, Quadrant 2

29. "Zc*" Extension for Code Size Reduction, Version 1.0.0

29.1. Zc* Overview

Zc* is a group of extensions that define subsets of the existing C extension (Zca, Zcd, Zcf) and new extensions which only contain 16-bit encodings.

Zcm* all reuse the encodings for c.fld, c.fsd, c.fldsp, c.fsdsp.

Table 24. Zc* extension overview
Instruction Zca Zcf Zcd Zcb Zcmp Zcmt

The Zca extension is added as way to refer to instructions in the C extension that do not include the floating-point loads and stores

C excl. c.f*

yes

The Zcf extension is added as a way to refer to compressed single-precision floating-point load/stores

c.flw

rv32

c.flwsp

rv32

c.fsw

rv32

c.fswsp

rv32

The Zcd extension is added as a way to refer to compressed double-precision floating-point load/stores

c.fld

yes

c.fldsp

yes

c.fsd

yes

c.fsdsp

yes

Simple operations for use on all architectures

c.lbu

yes

c.lh

yes

c.lhu

yes

c.sb

yes

c.sh

yes

c.zext.b

yes

c.sext.b

yes

c.zext.h

yes

c.sext.h

yes

c.zext.w

yes

c.mul

yes

c.not

yes

PUSH/POP and double move which overlap with c.fsdsp. Complex operations intended for embedded CPUs

cm.push

yes

cm.pop

yes

cm.popret

yes

cm.popretz

yes

cm.mva01s

yes

cm.mvsa01

yes

Table jump which overlaps with c.fsdsp. Complex operations intended for embedded CPUs

cm.jt

yes

cm.jalt

yes

29.2. C

The C extension is the superset of the following extensions:

  • Zca

  • Zcf if F is specified (RV32 only)

  • Zcd if D is specified

As C defines the same instructions as Zca, Zcf and Zcd, the rule is that:

  • C always implies Zca

  • C+F implies Zcf (RV32 only)

  • C+D implies Zcd

29.3. Zce

The Zce extension is intended to be used for microcontrollers, and includes all relevant Zc extensions.

  • Specifying Zce on RV32 without F includes Zca, Zcb, Zcmp, Zcmt

  • Specifying Zce on RV32 with F includes Zca, Zcb, Zcmp, Zcmt and Zcf

  • Specifying Zce on RV64 always includes Zca, Zcb, Zcmp, Zcmt

    • Zcf doesn’t exist for RV64

Therefore common ISA strings can be updated as follows to include the relevant Zc extensions, for example:

  • RV32IMC becomes RV32IM_Zce

  • RV32IMCF becomes RV32IMF_Zce

29.4. MISA.C

MISA.C is set if the following extensions are selected:

  • Zca and not F

  • Zca, Zcf and F is specified (RV32 only)

  • Zca, Zcf and Zcd if D is specified (RV32 only)

    • this configuration excludes Zcmp, Zcmt

  • Zca, Zcd if D is specified (RV64 only)

    • this configuration excludes Zcmp, Zcmt

29.5. Zca

The Zca extension is added as way to refer to instructions in the C extension that do not include the floating-point loads and stores.

Therefore it excluded all 16-bit floating point loads and stores: c.flw, c.flwsp, c.fsw, c.fswsp, c.fld, c.fldsp, c.fsd, c.fsdsp.

the C extension only includes F/D instructions when D and F are also specified

29.6. Zcf (RV32 only)

Zcf is the existing set of compressed single precision floating point loads and stores: c.flw, c.flwsp, c.fsw, c.fswsp.

Zcf is only relevant to RV32, it cannot be specified for RV64.

The Zcf extension depends on the Zca and F extensions.

29.7. Zcd

Zcd is the existing set of compressed double precision floating point loads and stores: c.fld, c.fldsp, c.fsd, c.fsdsp.

The Zcd extension depends on the Zca and D extensions.

29.8. Zcb

Zcb has simple code-size saving instructions which are easy to implement on all CPUs.

All encodings are currently reserved for all architectures, and have no conflicts with any existing extensions.

Zcb can be implemented on any CPU as the instructions are 16-bit versions of existing 32-bit instructions from the application class profile.

The Zcb extension depends on the Zca extension.

As shown on the individual instruction pages, many of the instructions in Zcb depend upon another extension being implemented. For example, c.mul is only implemented if M or Zmmul is implemented, and c.sext.b is only implemented if Zbb is implemented.

The c.mul encoding uses the CA register format along with other instructions such as c.sub, c.xor etc.

c.sext.w is a pseudoinstruction for c.addiw rd, 0 (RV64)
RV32 RV64 Mnemonic Instruction

yes

yes

c.lbu rd', uimm(rs1')

Load unsigned byte, 16-bit encoding

yes

yes

c.lhu rd', uimm(rs1')

Load unsigned halfword, 16-bit encoding

yes

yes

c.lh rd', uimm(rs1')

Load signed halfword, 16-bit encoding

yes

yes

c.sb rs2', uimm(rs1')

Store byte, 16-bit encoding

yes

yes

c.sh rs2', uimm(rs1')

Store halfword, 16-bit encoding

yes

yes

c.zext.b rsd'

Zero extend byte, 16-bit encoding

yes

yes

c.sext.b rsd'

Sign extend byte, 16-bit encoding

yes

yes

c.zext.h rsd'

Zero extend halfword, 16-bit encoding

yes

yes

c.sext.h rsd'

Sign extend halfword, 16-bit encoding

yes

c.zext.w rsd'

Zero extend word, 16-bit encoding

yes

yes

c.not rsd'

Bitwise not, 16-bit encoding

yes

yes

c.mul rsd', rs2'

Multiply, 16-bit encoding

29.9. Zcmp

The Zcmp extension is a set of instructions which may be executed as a series of existing 32-bit RISC-V instructions.

This extension reuses some encodings from c.fsdsp. Therefore it is incompatible with Zcd, which is included when C and D extensions are both present.

Zcmp is primarily targeted at embedded class CPUs due to implementation complexity. Additionally, it is not compatible with architecture class profiles.

The Zcmp extension depends on the Zca extension.

The PUSH/POP assembly syntax uses several variables, the meaning of which are:

  • reg_list is a list containing 1 to 13 registers (ra and 0 to 12 s registers)

    • valid values: {ra}, {ra, s0}, {ra, s0-s1}, {ra, s0-s2}, …​, {ra, s0-s8}, {ra, s0-s9}, {ra, s0-s11}

    • note that {ra, s0-s10} is not valid, giving 12 lists not 13 for better encoding

  • stack_adj is the total size of the stack frame.

    • valid values vary with register list length and the specific encoding, see the instruction pages for details.

RV32 RV64 Mnemonic Instruction

yes

yes

cm.push {reg_list}, -stack_adj

cm.push

yes

yes

cm.pop {reg_list}, stack_adj

cm.pop

yes

yes

cm.popret {reg_list}, stack_adj

cm.popret

yes

yes

cm.popretz {reg_list}, stack_adj

cm.popretz

yes

yes

cm.mva01s rs1', rs2'

Move two s0-s7 registers into a0-a1

yes

yes

cm.mvsa01 r1s', r2s'

Move a0-a1 into two different s0-s7 registers

29.10. Zcmt

Zcmt adds the table jump instructions and also adds the jvt CSR. The jvt CSR requires a state enable if Smstateen is implemented. See jvt CSR, table jump base vector and control register for details.

This extension reuses some encodings from c.fsdsp. Therefore it is incompatible with Zcd, which is included when C and D extensions are both present.

Zcmt is primarily targeted at embedded class CPUs due to implementation complexity. Additionally, it is not compatible with RVA profiles.

The Zcmt extension depends on the Zca and Zicsr extensions.

RV32 RV64 Mnemonic Instruction

yes

yes

cm.jt index

Jump via table

yes

yes

cm.jalt index

Jump and link via table

29.11. Zc instruction formats

Several instructions in this specification use the following new instruction formats.

Format instructions 15:10 9 8 7 6 5 4 3 2 1 0

CLB

c.lbu

funct6

rs1'

uimm

rd'

op

CSB

c.sb

funct6

rs1'

uimm

rs2'

op

CLH

c.lhu, c.lh

funct6

rs1'

funct1

uimm

rd'

op

CSH

c.sh

funct6

rs1'

funct1

uimm

rs2'

op

CU

c.[sz]ext.*, c.not

funct6

rd'/rs1'

funct5

op

CMMV

cm.mvsa01 cm.mva01s

funct6

r1s'

funct2

r2s'

op

CMJT

cm.jt cm.jalt

funct6

index

op

CMPP

cm.push*, cm.pop*

funct6

funct2

urlist

spimm

op

c.mul uses the existing CA format.

29.12. Zcb instructions

29.12.1. c.lbu

Synopsis:

Load unsigned byte, 16-bit encoding

Mnemonic:

c.lbu rd', uimm(rs1')

Encoding (RV32, RV64):

Diagram

The immediate offset is formed as follows:

  uimm[31:2] = 0;
  uimm[1]    = encoding[5];
  uimm[0]    = encoding[6];

Description:

This instruction loads a byte from the memory address formed by adding rs1' to the zero extended immediate uimm. The resulting byte is zero extended to XLEN bits and is written to rd'.

rd' and rs1' are from the standard 8-register set x8-x15.

Prerequisites:

None

Operation:

//This is not SAIL, it's pseudocode. The SAIL hasn't been written yet.

X(rdc) = EXTZ(mem[X(rs1c)+EXTZ(uimm)][7..0]);

29.12.2. c.lhu

Synopsis:

Load unsigned halfword, 16-bit encoding

Mnemonic:

c.lhu rd', uimm(rs1')

Encoding (RV32, RV64):

Diagram

The immediate offset is formed as follows:

  uimm[31:2] = 0;
  uimm[1]    = encoding[5];
  uimm[0]    = 0;

Description:

This instruction loads a halfword from the memory address formed by adding rs1' to the zero extended immediate uimm. The resulting halfword is zero extended to XLEN bits and is written to rd'.

rd' and rs1' are from the standard 8-register set x8-x15.

Prerequisites:

None

Operation:

//This is not SAIL, it's pseudocode. The SAIL hasn't been written yet.

X(rdc) = EXTZ(load_mem[X(rs1c)+EXTZ(uimm)][15..0]);

29.12.3. c.lh

Synopsis:

Load signed halfword, 16-bit encoding

Mnemonic:

c.lh rd', uimm(rs1')

Encoding (RV32, RV64):

Diagram

The immediate offset is formed as follows:

  uimm[31:2] = 0;
  uimm[1]    = encoding[5];
  uimm[0]    = 0;

Description:

This instruction loads a halfword from the memory address formed by adding rs1' to the zero extended immediate uimm. The resulting halfword is sign extended to XLEN bits and is written to rd'.

rd' and rs1' are from the standard 8-register set x8-x15.

Prerequisites:

None

Operation:

//This is not SAIL, it's pseudocode. The SAIL hasn't been written yet.

X(rdc) = EXTS(load_mem[X(rs1c)+EXTZ(uimm)][15..0]);

29.12.4. c.sb

Synopsis:

Store byte, 16-bit encoding

Mnemonic:

c.sb rs2', uimm(rs1')

Encoding (RV32, RV64):

Diagram

The immediate offset is formed as follows:

  uimm[31:2] = 0;
  uimm[1]    = encoding[5];
  uimm[0]    = encoding[6];

Description:

This instruction stores the least significant byte of rs2' to the memory address formed by adding rs1' to the zero extended immediate uimm.

rs1' and rs2' are from the standard 8-register set x8-x15.

Prerequisites:

None

Operation:

//This is not SAIL, it's pseudocode. The SAIL hasn't been written yet.

mem[X(rs1c)+EXTZ(uimm)][7..0] = X(rs2c)

29.12.5. c.sh

Synopsis:

Store halfword, 16-bit encoding

Mnemonic:

c.sh rs2', uimm(rs1')

Encoding (RV32, RV64):

Diagram

The immediate offset is formed as follows:

  uimm[31:2] = 0;
  uimm[1]    = encoding[5];
  uimm[0]    = 0;

Description:

This instruction stores the least significant halfword of rs2' to the memory address formed by adding rs1' to the zero extended immediate uimm.

rs1' and rs2' are from the standard 8-register set x8-x15.

Prerequisites:

None

Operation:

//This is not SAIL, it's pseudocode. The SAIL hasn't been written yet.

mem[X(rs1c)+EXTZ(uimm)][15..0] = X(rs2c)

29.12.6. c.zext.b

Synopsis:

Zero extend byte, 16-bit encoding

Mnemonic:

c.zext.b rd'/rs1'

Encoding (RV32, RV64):

Diagram

Description:

This instruction takes a single source/destination operand. It zero-extends the least-significant byte of the operand to XLEN bits by inserting zeros into all of the bits more significant than 7.

rd'/rs1' is from the standard 8-register set x8-x15.

Prerequisites:

None

32-bit equivalent:

andi rd'/rs1', rd'/rs1', 0xff

The SAIL module variable for rd'/rs1' is called rsdc.

Operation:

X(rsdc) = EXTZ(X(rsdc)[7..0]);

29.12.7. c.sext.b

Synopsis:

Sign extend byte, 16-bit encoding

Mnemonic:

c.sext.b rd'/rs1'

Encoding (RV32, RV64):

Diagram

Description:

This instruction takes a single source/destination operand. It sign-extends the least-significant byte in the operand to XLEN bits by copying the most-significant bit in the byte (i.e., bit 7) to all of the more-significant bits.

rd'/rs1' is from the standard 8-register set x8-x15.

Prerequisites:

Zbb is also required.

The SAIL module variable for rd'/rs1' is called rsdc.

Operation:

X(rsdc) = EXTS(X(rsdc)[7..0]);

29.12.8. c.zext.h

Synopsis:

Zero extend halfword, 16-bit encoding

Mnemonic:

c.zext.h rd'/rs1'

Encoding (RV32, RV64):

Diagram

Description:

This instruction takes a single source/destination operand. It zero-extends the least-significant halfword of the operand to XLEN bits by inserting zeros into all of the bits more significant than 15.

rd'/rs1' is from the standard 8-register set x8-x15.

Prerequisites:

Zbb is also required.

The SAIL module variable for rd'/rs1' is called rsdc.

Operation:

X(rsdc) = EXTZ(X(rsdc)[15..0]);

29.12.9. c.sext.h

Synopsis:

Sign extend halfword, 16-bit encoding

Mnemonic:

c.sext.h rd'/rs1'

Encoding (RV32, RV64):

Diagram

Description:

This instruction takes a single source/destination operand. It sign-extends the least-significant halfword in the operand to XLEN bits by copying the most-significant bit in the halfword (i.e., bit 15) to all of the more-significant bits.

rd'/rs1' is from the standard 8-register set x8-x15.

Prerequisites:

Zbb is also required.

The SAIL module variable for rd'/rs1' is called rsdc.

Operation:

X(rsdc) = EXTS(X(rsdc)[15..0]);

29.12.10. c.zext.w

Synopsis:

Zero extend word, 16-bit encoding

Mnemonic:

c.zext.w rd'/rs1'

Encoding (RV64):

Diagram

Description:

This instruction takes a single source/destination operand. It zero-extends the least-significant word of the operand to XLEN bits by inserting zeros into all of the bits more significant than 31.

rd'/rs1' is from the standard 8-register set x8-x15.

Prerequisites:

Zba is also required.

32-bit equivalent:

add.uw rd'/rs1', rd'/rs1', zero

The SAIL module variable for rd'/rs1' is called rsdc.

Operation:

X(rsdc) = EXTZ(X(rsdc)[31..0]);

29.12.11. c.not

Synopsis:

Bitwise not, 16-bit encoding

Mnemonic:

c.not rd'/rs1'

Encoding (RV32, RV64):

Diagram

Description:

This instruction takes the one’s complement of rd'/rs1' and writes the result to the same register.

rd'/rs1' is from the standard 8-register set x8-x15.

Prerequisites:

None

32-bit equivalent:

xori rd'/rs1', rd'/rs1', -1

The SAIL module variable for rd'/rs1' is called rsdc.

Operation:

X(rsdc) = X(rsdc) XOR -1;

29.12.12. c.mul

Synopsis:

Multiply, 16-bit encoding

Mnemonic:

c.mul rsd', rs2'

Encoding (RV32, RV64):

Diagram

Description:

This instruction multiplies XLEN bits of the source operands from rsd' and rs2' and writes the lowest XLEN bits of the result to rsd'.

rd'/rs1' and rs2' are from the standard 8-register set x8-x15.

Prerequisites:

M or Zmmul must be configured.

The SAIL module variable for rd'/rs1' is called rsdc, and for rs2' is called rs2c.

Operation:

let result_wide = to_bits(2 * sizeof(xlen), signed(X(rsdc)) * signed(X(rs2c)));
X(rsdc) = result_wide[(sizeof(xlen) - 1) .. 0];

29.13. PUSH/POP register instructions

These instructions are collectively referred to as PUSH/POP:

The term PUSH refers to cm.push.

The term POP refers to cm.pop.

The term POPRET refers to cm.popret and cm.popretz.

Common details for these instructions are in this section.

29.13.1. PUSH/POP functional overview

PUSH, POP, POPRET are used to reduce the size of function prologues and epilogues.

  1. The PUSH instruction

    • adjusts the stack pointer to create the stack frame

    • pushes (stores) the registers specified in the register list to the stack frame

  2. The POP instruction

    • pops (loads) the registers in the register list from the stack frame

    • adjusts the stack pointer to destroy the stack frame

  3. The POPRET instructions

    • pop (load) the registers in the register list from the stack frame

    • cm.popretz also moves zero into a0 as the return value

    • adjust the stack pointer to destroy the stack frame

    • execute a ret instruction to return from the function

29.13.2. Example usage

This example gives an illustration of the use of PUSH and POPRET.

The function processMarkers in the EMBench benchmark picojpeg in the following file on github: libpicojpeg.c

The prologue and epilogue compile with GCC10 to:

   0001098a <processMarkers>:
   1098a:       711d                    addi    sp,sp,-96 ;#cm.push(1)
   1098c:       c8ca                    sw      s2,80(sp) ;#cm.push(2)
   1098e:       c6ce                    sw      s3,76(sp) ;#cm.push(3)
   10990:       c4d2                    sw      s4,72(sp) ;#cm.push(4)
   10992:       ce86                    sw      ra,92(sp) ;#cm.push(5)
   10994:       cca2                    sw      s0,88(sp) ;#cm.push(6)
   10996:       caa6                    sw      s1,84(sp) ;#cm.push(7)
   10998:       c2d6                    sw      s5,68(sp) ;#cm.push(8)
   1099a:       c0da                    sw      s6,64(sp) ;#cm.push(9)
   1099c:       de5e                    sw      s7,60(sp) ;#cm.push(10)
   1099e:       dc62                    sw      s8,56(sp) ;#cm.push(11)
   109a0:       da66                    sw      s9,52(sp) ;#cm.push(12)
   109a2:       d86a                    sw      s10,48(sp);#cm.push(13)
   109a4:       d66e                    sw      s11,44(sp);#cm.push(14)
...
   109f4:       4501                    li      a0,0      ;#cm.popretz(1)
   109f6:       40f6                    lw      ra,92(sp) ;#cm.popretz(2)
   109f8:       4466                    lw      s0,88(sp) ;#cm.popretz(3)
   109fa:       44d6                    lw      s1,84(sp) ;#cm.popretz(4)
   109fc:       4946                    lw      s2,80(sp) ;#cm.popretz(5)
   109fe:       49b6                    lw      s3,76(sp) ;#cm.popretz(6)
   10a00:       4a26                    lw      s4,72(sp) ;#cm.popretz(7)
   10a02:       4a96                    lw      s5,68(sp) ;#cm.popretz(8)
   10a04:       4b06                    lw      s6,64(sp) ;#cm.popretz(9)
   10a06:       5bf2                    lw      s7,60(sp) ;#cm.popretz(10)
   10a08:       5c62                    lw      s8,56(sp) ;#cm.popretz(11)
   10a0a:       5cd2                    lw      s9,52(sp) ;#cm.popretz(12)
   10a0c:       5d42                    lw      s10,48(sp);#cm.popretz(13)
   10a0e:       5db2                    lw      s11,44(sp);#cm.popretz(14)
   10a10:       6125                    addi    sp,sp,96  ;#cm.popretz(15)
   10a12:       8082                    ret               ;#cm.popretz(16)

with the GCC option -msave-restore the output is the following:

0001080e <processMarkers>:
   1080e:       73a012ef                jal     t0,11f48 <__riscv_save_12>
   10812:       1101                    addi    sp,sp,-32
...
   10862:       4501                    li      a0,0
   10864:       6105                    addi    sp,sp,32
   10866:       71e0106f                j       11f84 <__riscv_restore_12>

with PUSH/POPRET this reduces to

0001080e <processMarkers>:
   1080e:       b8fa                    cm.push    {ra,s0-s11},-96
...
   10866:       bcfa                    cm.popretz {ra,s0-s11}, 96

The prologue / epilogue reduce from 60-bytes in the original code, to 14-bytes with -msave-restore, and to 4-bytes with PUSH and POPRET. As well as reducing the code-size PUSH and POPRET eliminate the branches from calling the millicode save/restore routines and so may also perform better.

The calls to <riscv_save_0>/<riscv_restore_0> become 64-bit when the target functions are out of the ±1MB range, increasing the prologue/epilogue size to 22-bytes.

POP is typically used in tail-calling sequences where ret is not used to return to ra after destroying the stack frame.
29.13.2.1. Stack pointer adjustment handling

The instructions all automatically adjust the stack pointer by enough to cover the memory required for the registers being saved or restored. Additionally the spimm field in the encoding allows the stack pointer to be adjusted in additional increments of 16-bytes. There is only a small restricted range available in the encoding; if the range is insufficient then a separate c.addi16sp can be used to increase the range.

29.13.2.2. Register list handling

There is no support for the {ra, s0-s10} register list without also adding s11. Therefore the {ra, s0-s11} register list must be used in this case.

29.13.3. PUSH/POP Fault handling

Correct execution requires that sp refers to idempotent memory (also see Non-idempotent memory handling), because the core must be able to handle traps detected during the sequence. The entire PUSH/POP sequence is re-executed after returning from the trap handler, and multiple traps are possible during the sequence.

If a trap occurs during the sequence then xEPC is updated with the PC of the instruction, xTVAL (if not read-only-zero) updated with the bad address if it was an access fault and xCAUSE updated with the type of trap.

It is implementation defined whether interrupts can also be taken during the sequence execution.

29.13.4. Software view of execution

29.13.4.1. Software view of the PUSH sequence

From a software perspective the PUSH sequence appears as:

  • A sequence of stores writing the bytes required by the pseudocode

    • The bytes may be written in any order.

    • The bytes may be grouped into larger accesses.

    • Any of the bytes may be written multiple times.

  • A stack pointer adjustment

If an implementation allows interrupts during the sequence, and the interrupt handler uses sp to allocate stack memory, then any stores which were executed before the interrupt may be overwritten by the handler. This is safe because the memory is idempotent and the stores will be re-executed when execution resumes.

The stack pointer adjustment must only be committed only when it is certain that the entire PUSH instruction will commit.

Stores may also return imprecise faults from the bus. It is platform defined whether the core implementation waits for the bus responses before continuing to the final stage of the sequence, or handles errors responses after completing the PUSH instruction.

For example:

cm.push  {ra, s0-s5}, -64

Appears to software as:

# any bytes from sp-1 to sp-28 may be written multiple times before
# the instruction completes therefore these updates may be visible in
# the interrupt/exception handler below the stack pointer
sw  s5, -4(sp)
sw  s4, -8(sp)
sw  s3,-12(sp)
sw  s2,-16(sp)
sw  s1,-20(sp)
sw  s0,-24(sp)
sw  ra,-28(sp)

# this must only execute once, and will only execute after all stores
# completed without any precise faults, therefore this update is only
# visible in the interrupt/exception handler if cm.push has completed
addi sp, sp, -64
29.13.4.2. Software view of the POP/POPRET sequence

From a software perspective the POP/POPRET sequence appears as:

  • A sequence of loads reading the bytes required by the pseudocode.

    • The bytes may be loaded in any order.

    • The bytes may be grouped into larger accesses.

    • Any of the bytes may be loaded multiple times.

  • A stack pointer adjustment

  • An optional li a0, 0

  • An optional ret

If a trap occurs during the sequence, then any loads which were executed before the trap may update architectural state. The loads will be re-executed once the trap handler completes, so the values will be overwritten. Therefore it is permitted for an implementation to update some of the destination registers before taking a fault.

The optional li a0, 0, stack pointer adjustment and optional ret must only be committed only when it is certain that the entire POP/POPRET instruction will commit.

For POPRET once the stack pointer adjustment has been committed the ret must execute.

For example:

cm.popretz {ra, s0-s3}, 32;

Appears to software as:

# any or all of these load instructions may execute multiple times
# therefore these updates may be visible in the interrupt/exception handler
lw   s3, 28(sp)
lw   s2, 24(sp)
lw   s1, 20(sp)
lw   s0, 16(sp)
lw   ra, 12(sp)

# these must only execute once, will only execute after all loads
# complete successfully all instructions must execute atomically
# therefore these updates are not visible in the interrupt/exception handler
li a0, 0
addi sp, sp, 32
ret

29.13.5. Non-idempotent memory handling

An implementation may have a requirement to issue a PUSH/POP instruction to non-idempotent memory.

If the core implementation does not support PUSH/POP to non-idempotent memories, the core may use an idempotency PMA to detect it and take a load (POP/POPRET) or store (PUSH) access fault exception in order to avoid unpredictable results.

Software should only use these instructions on non-idempotent memory regions when software can tolerate the required memory accesses being issued repeatedly in the case that they cause exceptions.

29.13.6. Example RV32I PUSH/POP sequences

The examples are included show the load/store series expansion and the stack adjustment. Examples of cm.popret and cm.popretz are not included, as the difference in the expanded sequence from cm.pop is trivial in all cases.

29.13.6.1. cm.push {ra, s0-s2}, -64

Encoding: rlist=7, spimm=3

expands to:

sw  s2,  -4(sp);
sw  s1,  -8(sp);
sw  s0, -12(sp);
sw  ra, -16(sp);
addi sp, sp, -64;
29.13.6.2. cm.push {ra, s0-s11}, -112

Encoding: rlist=15, spimm=3

expands to:

sw  s11,  -4(sp);
sw  s10,  -8(sp);
sw  s9,  -12(sp);
sw  s8,  -16(sp);
sw  s7,  -20(sp);
sw  s6,  -24(sp);
sw  s5,  -28(sp);
sw  s4,  -32(sp);
sw  s3,  -36(sp);
sw  s2,  -40(sp);
sw  s1,  -44(sp);
sw  s0,  -48(sp);
sw  ra,  -52(sp);
addi sp, sp, -112;
29.13.6.3. cm.pop {ra}, 16

Encoding: rlist=4, spimm=0

expands to:

lw   ra, 12(sp);
addi sp, sp, 16;
29.13.6.4. cm.pop {ra, s0-s3}, 48

Encoding: rlist=8, spimm=1

expands to:

lw   s3, 44(sp);
lw   s2, 40(sp);
lw   s1, 36(sp);
lw   s0, 32(sp);
lw   ra, 28(sp);
addi sp, sp, 48;
29.13.6.5. cm.pop {ra, s0-s4}, 64

Encoding: rlist=9, spimm=2

expands to:

lw   s4, 60(sp);
lw   s3, 56(sp);
lw   s2, 52(sp);
lw   s1, 48(sp);
lw   s0, 44(sp);
lw   ra, 40(sp);
addi sp, sp, 64;

29.13.7. cm.push

Synopsis:

Create stack frame: store ra and 0 to 12 saved registers to the stack frame, optionally allocate additional stack space.

Mnemonic:

cm.push {reg_list}, -stack_adj

Encoding (RV32, RV64):

Diagram

rlist values 0 to 3 are reserved for a future EABI variant called cm.push.e

Assembly Syntax:

cm.push {reg_list},  -stack_adj
cm.push {xreg_list}, -stack_adj

The variables used in the assembly syntax are defined below.

RV32E:

switch (rlist){
  case  4: {reg_list="ra";         xreg_list="x1";}
  case  5: {reg_list="ra, s0";     xreg_list="x1, x8";}
  case  6: {reg_list="ra, s0-s1";  xreg_list="x1, x8-x9";}
  default: reserved();
}
stack_adj      = stack_adj_base + spimm[5:4] * 16;
RV32I, RV64:

switch (rlist){
  case  4: {reg_list="ra";         xreg_list="x1";}
  case  5: {reg_list="ra, s0";     xreg_list="x1, x8";}
  case  6: {reg_list="ra, s0-s1";  xreg_list="x1, x8-x9";}
  case  7: {reg_list="ra, s0-s2";  xreg_list="x1, x8-x9, x18";}
  case  8: {reg_list="ra, s0-s3";  xreg_list="x1, x8-x9, x18-x19";}
  case  9: {reg_list="ra, s0-s4";  xreg_list="x1, x8-x9, x18-x20";}
  case 10: {reg_list="ra, s0-s5";  xreg_list="x1, x8-x9, x18-x21";}
  case 11: {reg_list="ra, s0-s6";  xreg_list="x1, x8-x9, x18-x22";}
  case 12: {reg_list="ra, s0-s7";  xreg_list="x1, x8-x9, x18-x23";}
  case 13: {reg_list="ra, s0-s8";  xreg_list="x1, x8-x9, x18-x24";}
  case 14: {reg_list="ra, s0-s9";  xreg_list="x1, x8-x9, x18-x25";}
  //note - to include s10, s11 must also be included
  case 15: {reg_list="ra, s0-s11"; xreg_list="x1, x8-x9, x18-x27";}
  default: reserved();
}
stack_adj      = stack_adj_base + spimm[5:4] * 16;
RV32E:

stack_adj_base = 16;
Valid values:
stack_adj      = [16|32|48|64];
RV32I:

switch (rlist) {
  case  4.. 7: stack_adj_base = 16;
  case  8..11: stack_adj_base = 32;
  case 12..14: stack_adj_base = 48;
  case     15: stack_adj_base = 64;
}

Valid values:
switch (rlist) {
  case  4.. 7: stack_adj = [16|32|48| 64];
  case  8..11: stack_adj = [32|48|64| 80];
  case 12..14: stack_adj = [48|64|80| 96];
  case     15: stack_adj = [64|80|96|112];
}
RV64:

switch (rlist) {
  case  4.. 5: stack_adj_base =  16;
  case  6.. 7: stack_adj_base =  32;
  case  8.. 9: stack_adj_base =  48;
  case 10..11: stack_adj_base =  64;
  case 12..13: stack_adj_base =  80;
  case     14: stack_adj_base =  96;
  case     15: stack_adj_base = 112;
}

Valid values:
switch (rlist) {
  case  4.. 5: stack_adj = [ 16| 32| 48| 64];
  case  6.. 7: stack_adj = [ 32| 48| 64| 80];
  case  8.. 9: stack_adj = [ 48| 64| 80| 96];
  case 10..11: stack_adj = [ 64| 80| 96|112];
  case 12..13: stack_adj = [ 80| 96|112|128];
  case     14: stack_adj = [ 96|112|128|144];
  case     15: stack_adj = [112|128|144|160];
}

Description:

This instruction pushes (stores) the registers in reg_list to the memory below the stack pointer, and then creates the stack frame by decrementing the stack pointer by stack_adj, including any additional stack space requested by the value of spimm.

All ABI register mappings are for the UABI. An EABI version is planned once the EABI is frozen.

For further information see PUSH/POP Register Instructions.

Stack Adjustment Calculation:

stack_adj_base is the minimum number of bytes, in multiples of 16-byte address increments, required to cover the registers in the list.

spimm is the number of additional 16-byte address increments allocated for the stack frame.

The total stack adjustment represents the total size of the stack frame, which is stack_adj_base added to spimm scaled by 16, as defined above.

Prerequisites:

None

32-bit equivalent:

No direct equivalent encoding exists

Operation:

The first section of pseudocode may be executed multiple times before the instruction successfully completes.

//This is not SAIL, it's pseudocode. The SAIL hasn't been written yet.

if (XLEN==32) bytes=4; else bytes=8;

addr=sp-bytes;
for(i in 27,26,25,24,23,22,21,20,19,18,9,8,1)  {
  //if register i is in xreg_list
  if (xreg_list[i]) {
    switch(bytes) {
      4:  asm("sw x[i], 0(addr)");
      8:  asm("sd x[i], 0(addr)");
    }
    addr-=bytes;
  }
}

The final section of pseudocode executes atomically, and only executes if the section above completes without any exceptions or interrupts.

//This is not SAIL, it's pseudocode. The SAIL hasn't been written yet.

sp-=stack_adj;

29.13.8. cm.pop

Synopsis:

Destroy stack frame: load ra and 0 to 12 saved registers from the stack frame, deallocate the stack frame.

Mnemonic:

cm.pop {reg_list}, stack_adj

Encoding (RV32, RV64):

Diagram

rlist values 0 to 3 are reserved for a future EABI variant called cm.pop.e

Assembly Syntax:

cm.pop {reg_list},  stack_adj
cm.pop {xreg_list}, stack_adj

The variables used in the assembly syntax are defined below.

RV32E:

switch (rlist){
  case  4: {reg_list="ra";         xreg_list="x1";}
  case  5: {reg_list="ra, s0";     xreg_list="x1, x8";}
  case  6: {reg_list="ra, s0-s1";  xreg_list="x1, x8-x9";}
  default: reserved();
}
stack_adj      = stack_adj_base + spimm[5:4] * 16;
RV32I, RV64:

switch (rlist){
  case  4: {reg_list="ra";         xreg_list="x1";}
  case  5: {reg_list="ra, s0";     xreg_list="x1, x8";}
  case  6: {reg_list="ra, s0-s1";  xreg_list="x1, x8-x9";}
  case  7: {reg_list="ra, s0-s2";  xreg_list="x1, x8-x9, x18";}
  case  8: {reg_list="ra, s0-s3";  xreg_list="x1, x8-x9, x18-x19";}
  case  9: {reg_list="ra, s0-s4";  xreg_list="x1, x8-x9, x18-x20";}
  case 10: {reg_list="ra, s0-s5";  xreg_list="x1, x8-x9, x18-x21";}
  case 11: {reg_list="ra, s0-s6";  xreg_list="x1, x8-x9, x18-x22";}
  case 12: {reg_list="ra, s0-s7";  xreg_list="x1, x8-x9, x18-x23";}
  case 13: {reg_list="ra, s0-s8";  xreg_list="x1, x8-x9, x18-x24";}
  case 14: {reg_list="ra, s0-s9";  xreg_list="x1, x8-x9, x18-x25";}
  //note - to include s10, s11 must also be included
  case 15: {reg_list="ra, s0-s11"; xreg_list="x1, x8-x9, x18-x27";}
  default: reserved();
}
stack_adj      = stack_adj_base + spimm[5:4] * 16;
RV32E:

stack_adj_base = 16;
Valid values:
stack_adj      = [16|32|48|64];
RV32I:

switch (rlist) {
  case  4.. 7: stack_adj_base = 16;
  case  8..11: stack_adj_base = 32;
  case 12..14: stack_adj_base = 48;
  case     15: stack_adj_base = 64;
}

Valid values:
switch (rlist) {
  case  4.. 7: stack_adj = [16|32|48| 64];
  case  8..11: stack_adj = [32|48|64| 80];
  case 12..14: stack_adj = [48|64|80| 96];
  case     15: stack_adj = [64|80|96|112];
}
RV64:

switch (rlist) {
  case  4.. 5: stack_adj_base =  16;
  case  6.. 7: stack_adj_base =  32;
  case  8.. 9: stack_adj_base =  48;
  case 10..11: stack_adj_base =  64;
  case 12..13: stack_adj_base =  80;
  case     14: stack_adj_base =  96;
  case     15: stack_adj_base = 112;
}

Valid values:
switch (rlist) {
  case  4.. 5: stack_adj = [ 16| 32| 48| 64];
  case  6.. 7: stack_adj = [ 32| 48| 64| 80];
  case  8.. 9: stack_adj = [ 48| 64| 80| 96];
  case 10..11: stack_adj = [ 64| 80| 96|112];
  case 12..13: stack_adj = [ 80| 96|112|128];
  case     14: stack_adj = [ 96|112|128|144];
  case     15: stack_adj = [112|128|144|160];
}

Description:

This instruction pops (loads) the registers in reg_list from stack memory, and then adjusts the stack pointer by stack_adj.

All ABI register mappings are for the UABI. An EABI version is planned once the EABI is frozen.

For further information see PUSH/POP Register Instructions.

Stack Adjustment Calculation:

stack_adj_base is the minimum number of bytes, in multiples of 16-byte address increments, required to cover the registers in the list.

spimm is the number of additional 16-byte address increments allocated for the stack frame.

The total stack adjustment represents the total size of the stack frame, which is stack_adj_base added to spimm scaled by 16, as defined above.

Prerequisites:

None

32-bit equivalent:

No direct equivalent encoding exists

Operation:

The first section of pseudocode may be executed multiple times before the instruction successfully completes.

//This is not SAIL, it's pseudocode. The SAIL hasn't been written yet.

if (XLEN==32) bytes=4; else bytes=8;

addr=sp+stack_adj-bytes;
for(i in 27,26,25,24,23,22,21,20,19,18,9,8,1)  {
  //if register i is in xreg_list
  if (xreg_list[i]) {
    switch(bytes) {
      4:  asm("lw x[i], 0(addr)");
      8:  asm("ld x[i], 0(addr)");
    }
    addr-=bytes;
  }
}

The final section of pseudocode executes atomically, and only executes if the section above completes without any exceptions or interrupts.

//This is not SAIL, it's pseudocode. The SAIL hasn't been written yet.

sp+=stack_adj;

29.13.9. cm.popretz

Synopsis:

Destroy stack frame: load ra and 0 to 12 saved registers from the stack frame, deallocate the stack frame, move zero into a0, return to ra.

Mnemonic:

cm.popretz {reg_list}, stack_adj

Encoding (RV32, RV64):

Diagram

rlist values 0 to 3 are reserved for a future EABI variant called cm.popretz.e

Assembly Syntax:

cm.popretz {reg_list},  stack_adj
cm.popretz {xreg_list}, stack_adj
RV32E:

switch (rlist){
  case  4: {reg_list="ra";         xreg_list="x1";}
  case  5: {reg_list="ra, s0";     xreg_list="x1, x8";}
  case  6: {reg_list="ra, s0-s1";  xreg_list="x1, x8-x9";}
  default: reserved();
}
stack_adj      = stack_adj_base + spimm[5:4] * 16;
RV32I, RV64:

switch (rlist){
  case  4: {reg_list="ra";         xreg_list="x1";}
  case  5: {reg_list="ra, s0";     xreg_list="x1, x8";}
  case  6: {reg_list="ra, s0-s1";  xreg_list="x1, x8-x9";}
  case  7: {reg_list="ra, s0-s2";  xreg_list="x1, x8-x9, x18";}
  case  8: {reg_list="ra, s0-s3";  xreg_list="x1, x8-x9, x18-x19";}
  case  9: {reg_list="ra, s0-s4";  xreg_list="x1, x8-x9, x18-x20";}
  case 10: {reg_list="ra, s0-s5";  xreg_list="x1, x8-x9, x18-x21";}
  case 11: {reg_list="ra, s0-s6";  xreg_list="x1, x8-x9, x18-x22";}
  case 12: {reg_list="ra, s0-s7";  xreg_list="x1, x8-x9, x18-x23";}
  case 13: {reg_list="ra, s0-s8";  xreg_list="x1, x8-x9, x18-x24";}
  case 14: {reg_list="ra, s0-s9";  xreg_list="x1, x8-x9, x18-x25";}
  //note - to include s10, s11 must also be included
  case 15: {reg_list="ra, s0-s11"; xreg_list="x1, x8-x9, x18-x27";}
  default: reserved();
}
stack_adj      = stack_adj_base + spimm[5:4] * 16;
RV32E:

stack_adj_base = 16;
Valid values:
stack_adj      = [16|32|48|64];
RV32I:

switch (rlist) {
  case  4.. 7: stack_adj_base = 16;
  case  8..11: stack_adj_base = 32;
  case 12..14: stack_adj_base = 48;
  case     15: stack_adj_base = 64;
}

Valid values:
switch (rlist) {
  case  4.. 7: stack_adj = [16|32|48| 64];
  case  8..11: stack_adj = [32|48|64| 80];
  case 12..14: stack_adj = [48|64|80| 96];
  case     15: stack_adj = [64|80|96|112];
}
RV64:

switch (rlist) {
  case  4.. 5: stack_adj_base =  16;
  case  6.. 7: stack_adj_base =  32;
  case  8.. 9: stack_adj_base =  48;
  case 10..11: stack_adj_base =  64;
  case 12..13: stack_adj_base =  80;
  case     14: stack_adj_base =  96;
  case     15: stack_adj_base = 112;
}

Valid values:
switch (rlist) {
  case  4.. 5: stack_adj = [ 16| 32| 48| 64];
  case  6.. 7: stack_adj = [ 32| 48| 64| 80];
  case  8.. 9: stack_adj = [ 48| 64| 80| 96];
  case 10..11: stack_adj = [ 64| 80| 96|112];
  case 12..13: stack_adj = [ 80| 96|112|128];
  case     14: stack_adj = [ 96|112|128|144];
  case     15: stack_adj = [112|128|144|160];
}

Description:

This instruction pops (loads) the registers in reg_list from stack memory, adjusts the stack pointer by stack_adj, moves zero into a0 and then returns to ra.

All ABI register mappings are for the UABI. An EABI version is planned once the EABI is frozen.

For further information see PUSH/POP Register Instructions.

Stack Adjustment Calculation:

stack_adj_base is the minimum number of bytes, in multiples of 16-byte address increments, required to cover the registers in the list.

spimm is the number of additional 16-byte address increments allocated for the stack frame.

The total stack adjustment represents the total size of the stack frame, which is stack_adj_base added to spimm scaled by 16, as defined above.

Prerequisites:

None

32-bit equivalent:

No direct equivalent encoding exists

Operation:

The first section of pseudocode may be executed multiple times before the instruction successfully completes.

//This is not SAIL, it's pseudocode. The SAIL hasn't been written yet.

if (XLEN==32) bytes=4; else bytes=8;

addr=sp+stack_adj-bytes;
for(i in 27,26,25,24,23,22,21,20,19,18,9,8,1)  {
  //if register i is in xreg_list
  if (xreg_list[i]) {
    switch(bytes) {
      4:  asm("lw x[i], 0(addr)");
      8:  asm("ld x[i], 0(addr)");
    }
    addr-=bytes;
  }
}

The final section of pseudocode executes atomically, and only executes if the section above completes without any exceptions or interrupts.

The li a0, 0 could be executed more than once, but is included in the atomic section for convenience.

 
//This is not SAIL, it's pseudocode. The SAIL hasn't been written yet.

asm("li a0, 0");
sp+=stack_adj;
asm("ret");

29.13.10. cm.popret

Synopsis:

Destroy stack frame: load ra and 0 to 12 saved registers from the stack frame, deallocate the stack frame, return to ra.

Mnemonic:

cm.popret {reg_list}, stack_adj

Encoding (RV32, RV64):

Diagram

rlist values 0 to 3 are reserved for a future EABI variant called cm.popret.e

Assembly Syntax:

cm.popret {reg_list},  stack_adj
cm.popret {xreg_list}, stack_adj

The variables used in the assembly syntax are defined below.

RV32E:

switch (rlist){
  case  4: {reg_list="ra";         xreg_list="x1";}
  case  5: {reg_list="ra, s0";     xreg_list="x1, x8";}
  case  6: {reg_list="ra, s0-s1";  xreg_list="x1, x8-x9";}
  default: reserved();
}
stack_adj      = stack_adj_base + spimm[5:4] * 16;
RV32I, RV64:

switch (rlist){
  case  4: {reg_list="ra";         xreg_list="x1";}
  case  5: {reg_list="ra, s0";     xreg_list="x1, x8";}
  case  6: {reg_list="ra, s0-s1";  xreg_list="x1, x8-x9";}
  case  7: {reg_list="ra, s0-s2";  xreg_list="x1, x8-x9, x18";}
  case  8: {reg_list="ra, s0-s3";  xreg_list="x1, x8-x9, x18-x19";}
  case  9: {reg_list="ra, s0-s4";  xreg_list="x1, x8-x9, x18-x20";}
  case 10: {reg_list="ra, s0-s5";  xreg_list="x1, x8-x9, x18-x21";}
  case 11: {reg_list="ra, s0-s6";  xreg_list="x1, x8-x9, x18-x22";}
  case 12: {reg_list="ra, s0-s7";  xreg_list="x1, x8-x9, x18-x23";}
  case 13: {reg_list="ra, s0-s8";  xreg_list="x1, x8-x9, x18-x24";}
  case 14: {reg_list="ra, s0-s9";  xreg_list="x1, x8-x9, x18-x25";}
  //note - to include s10, s11 must also be included
  case 15: {reg_list="ra, s0-s11"; xreg_list="x1, x8-x9, x18-x27";}
  default: reserved();
}
stack_adj      = stack_adj_base + spimm[5:4] * 16;
RV32E:

stack_adj_base = 16;
Valid values:
stack_adj      = [16|32|48|64];
RV32I:

switch (rlist) {
  case  4.. 7: stack_adj_base = 16;
  case  8..11: stack_adj_base = 32;
  case 12..14: stack_adj_base = 48;
  case     15: stack_adj_base = 64;
}

Valid values:
switch (rlist) {
  case  4.. 7: stack_adj = [16|32|48| 64];
  case  8..11: stack_adj = [32|48|64| 80];
  case 12..14: stack_adj = [48|64|80| 96];
  case     15: stack_adj = [64|80|96|112];
}
RV64:

switch (rlist) {
  case  4.. 5: stack_adj_base =  16;
  case  6.. 7: stack_adj_base =  32;
  case  8.. 9: stack_adj_base =  48;
  case 10..11: stack_adj_base =  64;
  case 12..13: stack_adj_base =  80;
  case     14: stack_adj_base =  96;
  case     15: stack_adj_base = 112;
}

Valid values:
switch (rlist) {
  case  4.. 5: stack_adj = [ 16| 32| 48| 64];
  case  6.. 7: stack_adj = [ 32| 48| 64| 80];
  case  8.. 9: stack_adj = [ 48| 64| 80| 96];
  case 10..11: stack_adj = [ 64| 80| 96|112];
  case 12..13: stack_adj = [ 80| 96|112|128];
  case     14: stack_adj = [ 96|112|128|144];
  case     15: stack_adj = [112|128|144|160];
}

Description:

This instruction pops (loads) the registers in reg_list from stack memory, adjusts the stack pointer by stack_adj and then returns to ra.

All ABI register mappings are for the UABI. An EABI version is planned once the EABI is frozen.

For further information see PUSH/POP Register Instructions.

Stack Adjustment Calculation:

stack_adj_base is the minimum number of bytes, in multiples of 16-byte address increments, required to cover the registers in the list.

spimm is the number of additional 16-byte address increments allocated for the stack frame.

The total stack adjustment represents the total size of the stack frame, which is stack_adj_base added to spimm scaled by 16, as defined above.

Prerequisites:

None

32-bit equivalent:

No direct equivalent encoding exists

Operation:

The first section of pseudocode may be executed multiple times before the instruction successfully completes.

//This is not SAIL, it's pseudocode. The SAIL hasn't been written yet.

if (XLEN==32) bytes=4; else bytes=8;

addr=sp+stack_adj-bytes;
for(i in 27,26,25,24,23,22,21,20,19,18,9,8,1)  {
  //if register i is in xreg_list
  if (xreg_list[i]) {
    switch(bytes) {
      4:  asm("lw x[i], 0(addr)");
      8:  asm("ld x[i], 0(addr)");
    }
    addr-=bytes;
  }
}

The final section of pseudocode executes atomically, and only executes if the section above completes without any exceptions or interrupts.

//This is not SAIL, it's pseudocode. The SAIL hasn't been written yet.

sp+=stack_adj;
asm("ret");

29.13.11. cm.mvsa01

Synopsis:

Move a0-a1 into two registers of s0-s7

Mnemonic:

cm.mvsa01 r1s', r2s'

Encoding (RV32, RV64):

Diagram

For the encoding to be legal r1s' != r2s'.

Assembly Syntax:

cm.mvsa01 r1s', r2s'

Description: This instruction moves a0 into r1s' and a1 into r2s'. r1s' and r2s' must be different. The execution is atomic, so it is not possible to observe state where only one of r1s' or r2s' has been updated.

The encoding uses sreg number specifiers instead of xreg number specifiers to save encoding space. The mapping between them is specified in the pseudocode below.

The s register mapping is taken from the UABI, and may not match the currently unratified EABI. cm.mvsa01.e may be included in the future.

Prerequisites:

None

32-bit equivalent:

No direct equivalent encoding exists.

Operation:

//This is not SAIL, it's pseudocode. The SAIL hasn't been written yet.
if (RV32E && (r1sc>1 || r2sc>1)) {
  reserved();
}
xreg1 = {r1sc[2:1]>0,r1sc[2:1]==0,r1sc[2:0]};
xreg2 = {r2sc[2:1]>0,r2sc[2:1]==0,r2sc[2:0]};
X[xreg1] = X[10];
X[xreg2] = X[11];

29.13.12. cm.mva01s

Synopsis:

Move two s0-s7 registers into a0-a1

Mnemonic:

cm.mva01s r1s', r2s'

Encoding (RV32, RV64):

Diagram

Assembly Syntax:

cm.mva01s r1s', r2s'

Description: This instruction moves r1s' into a0 and r2s' into a1. The execution is atomic, so it is not possible to observe state where only one of a0 or a1 have been updated.

The encoding uses sreg number specifiers instead of xreg number specifiers to save encoding space. The mapping between them is specified in the pseudocode below.

The s register mapping is taken from the UABI, and may not match the currently unratified EABI. cm.mva01s.e may be included in the future.

Prerequisites:

None

32-bit equivalent:

No direct equivalent encoding exists.

Operation:

//This is not SAIL, it's pseudocode. The SAIL hasn't been written yet.
if (RV32E && (r1sc>1 || r2sc>1)) {
  reserved();
}
xreg1 = {r1sc[2:1]>0,r1sc[2:1]==0,r1sc[2:0]};
xreg2 = {r2sc[2:1]>0,r2sc[2:1]==0,r2sc[2:0]};
X[10] = X[xreg1];
X[11] = X[xreg2];

29.14. Table Jump Overview

cm.jt (Jump via table) and cm.jalt (Jump and link via table) are referred to as table jump.

Table jump uses a 256-entry XLEN wide table in instruction memory to contain function addresses. The table must be a minimum of 64-byte aligned.

Table entries follow the current data endianness. This is different from normal instruction fetch which is always little-endian.

cm.jt and cm.jalt encodings index the table, giving access to functions within the full XLEN wide address space.

This is used as a form of dictionary compression to reduce the code size of jal / auipc+jalr / jr / auipc+jr instructions.

Table jump allows the linker to replace the following instruction sequences with a cm.jt or cm.jalt encoding, and an entry in the table:

  • 32-bit j calls

  • 32-bit jal ra calls

  • 64-bit auipc+jr calls to fixed locations

  • 64-bit auipc+jalr ra calls to fixed locations

    • The auipc+jr/jalr sequence is used because the offset from the PC is out of the ±1MB range.

If a return address stack is implemented, then as cm.jalt is equivalent to jal ra, it pushes to the stack.

29.14.1. jvt

The base of the table is in the jvt CSR (see jvt CSR, table jump base vector and control register), each table entry is XLEN bits.

If the same function is called with and without linking then it must have two entries in the table. This is typically caused by the same function being called with and without tail calling.

29.14.2. Table Jump Fault handling

For a table jump instruction, the table entry that the instruction selects is considered an extension of the instruction itself. Hence, the execution of a table jump instruction involves two instruction fetches, the first to read the instruction (cm.jt/cm.jalt) and the second to read from the jump vector table (JVT). Both instruction fetches are implicit reads, and both require execute permission; read permission is irrelevant. It is recommended that the second fetch be ignored for hardware triggers and breakpoints.

Memory writes to the jump vector table require an instruction barrier (fence.i) to guarantee that they are visible to the instruction fetch.

Multiple contexts may have different jump vector tables. JVT may be switched between them without an instruction barrier if the tables have not been updated in memory since the last fence.i.

If an exception occurs on either instruction fetch, xEPC is set to the PC of the table jump instruction, xCAUSE is set as expected for the type of fault and xTVAL (if not set to zero) contains the fetch address which caused the fault.

29.14.3. jvt CSR

Synopsis:

Table jump base vector and control register

Address:

0x0017

Permissions:

URW

Format (RV32):

Diagram

Format (RV64):

Diagram

Description:

The jvt register is an XLEN-bit WARL read/write register that holds the jump table configuration, consisting of the jump table base address (BASE) and the jump table mode (MODE).

If Section 29.10 is implemented then jvt must also be implemented, but can contain a read-only value. If jvt is writable, the set of values the register may hold can vary by implementation. The value in the BASE field must always be aligned on a 64-byte boundary.

jvt.base is a virtual address, whenever virtual memory is enabled.

The memory pointed to by jvt.base is treated as instruction memory for the purpose of executing table jump instructions, implying execute access permission.

Table 25. jvt.mode definition
jvt.mode Comment

000000

Jump table mode

others

reserved for future standard use

jvt.mode is a WARL field, so can only be programmed to modes which are implemented. Therefore the discovery mechanism is to attempt to program different modes and read back the values to see which are available. Jump table mode must be implemented.

in future the RISC-V Unified Discovery method will report the available modes.

Architectural State:

jvt CSR adds architectural state to the system software context (such as an OS process), therefore must be saved/restored on context switches.

State Enable:

If the Smstateen extension is implemented, then bit 2 in mstateen0, sstateen0, and hstateen0 is implemented. If bit 2 of a controlling stateen0 CSR is zero, then access to the jvt CSR and execution of a cm.jalt or cm.jt instruction by a lower privilege level results in an Illegal Instruction trap (or, if appropriate, a Virtual Instruction trap).

29.14.4. cm.jt

Synopsis:

jump via table

Mnemonic:

cm.jt index

Encoding (RV32, RV64):

Diagram

For this encoding to decode as cm.jt, index<32, otherwise it decodes as cm.jalt, see Jump and link via table.

If jvt.mode = 0 (Jump Table Mode) then cm.jt behaves as specified here. If jvt.mode is a reserved value, then cm.jt is also reserved. In the future other defined values of jvt.mode may change the behaviour of cm.jt.

Assembly Syntax:

cm.jt index

Description:

cm.jt reads an entry from the jump vector table in memory and jumps to the address that was read.

For further information see Table Jump Overview.

Prerequisites:

None

32-bit equivalent:

No direct equivalent encoding exists.

Operation:

//This is not SAIL, it's pseudocode. The SAIL hasn't been written yet.

# target_address is temporary internal state, it doesn't represent a real register
# InstMemory is byte indexed

switch(XLEN) {
  32:  table_address[XLEN-1:0] = jvt.base + (index<<2);
  64:  table_address[XLEN-1:0] = jvt.base + (index<<3);
}

//fetch from the jump table
target_address[XLEN-1:0] = InstMemory[table_address][XLEN-1:0];

j target_address[XLEN-1:0]&~0x1;

29.14.5. cm.jalt

Synopsis:

jump via table with optional link

Mnemonic:

cm.jalt index

Encoding (RV32, RV64):

Diagram

For this encoding to decode as cm.jalt, index>=32, otherwise it decodes as cm.jt, see Jump via table.

If jvt.mode = 0 (Jump Table Mode) then cm.jalt behaves as specified here. If jvt.mode is a reserved value, then cm.jalt is also reserved. In the future other defined values of jvt.mode may change the behaviour of cm.jalt.

Assembly Syntax:

cm.jalt index

Description:

cm.jalt reads an entry from the jump vector table in memory and jumps to the address that was read, linking to ra.

For further information see Table Jump Overview.

Prerequisites:

None

32-bit equivalent:

No direct equivalent encoding exists.

Operation:

//This is not SAIL, it's pseudocode. The SAIL hasn't been written yet.

# target_address is temporary internal state, it doesn't represent a real register
# InstMemory is byte indexed

switch(XLEN) {
  32:  table_address[XLEN-1:0] = jvt.base + (index<<2);
  64:  table_address[XLEN-1:0] = jvt.base + (index<<3);
}

//fetch from the jump table
target_address[XLEN-1:0] = InstMemory[table_address][XLEN-1:0];

jal ra, target_address[XLEN-1:0]&~0x1;

30. "B" Extension for Bit Manipulation, Version 1.0.0

The B standard extension comprises instructions provided by the Zba, Zbb, and Zbs extensions.

30.1. Zb* Overview

The bit-manipulation (bitmanip) extension collection is comprised of several component extensions to the base RISC-V architecture that are intended to provide some combination of code size reduction, performance improvement, and energy reduction. While the instructions are intended to have general use, some instructions are more useful in some domains than others. Hence, several smaller bitmanip extensions are provided. Each of these smaller extensions is grouped by common function and use case, and each has its own Zb*-extension name.

Each bitmanip extension includes a group of several bitmanip instructions that have similar purposes and that can often share the same logic. Some instructions are available in only one extension while others are available in several. The instructions have mnemonics and encodings that are independent of the extensions in which they appear. Thus, when implementing extensions with overlapping instructions, there is no redundancy in logic or encoding.

The bitmanip extensions are defined for RV32 and RV64. Most of the instructions are expected to be forward compatible with RV128. While the shift-immediate instructions are defined to have at most a 6-bit immediate field, a 7th bit is available in the encoding space should this be needed for RV128.

30.2. Word Instructions

The bitmanip extension follows the convention in RV64 that w-suffixed instructions (without a dot before the w) ignore the upper 32 bits of their inputs, operate on the least-significant 32-bits as signed values and produce a 32-bit signed result that is sign-extended to XLEN.

Bitmanip instructions with the suffix .uw have one operand that is an unsigned 32-bit value that is extracted from the least significant 32 bits of the specified register. Other than that, these perform full XLEN operations.

Bitmanip instructions with the suffix .b, .h and .w only look at the least significant 8-bits, 16-bits and 32-bits of the input (respectively) and produce an XLEN-wide result that is sign-extended or zero-extended, based on the specific instruction.

30.3. Pseudocode for instruction semantics

The semantics of each instruction in Instructions (in alphabetical order) is expressed in a SAIL-like syntax.

30.4. Extensions

The first group of bitmanip extensions to be released for Public Review are:

Below is a list of all of the instructions that are included in these extensions along with their specific mapping:

RV32 RV64 Mnemonic Instruction Zba Zbb Zbc Zbs

✓

add.uw rd, rs1, rs2

Add unsigned word

✓

✓

✓

andn rd, rs1, rs2

AND with inverted operand

✓

✓

✓

clmul rd, rs1, rs2

Carry-less multiply (low-part)

✓

✓

✓

clmulh rd, rs1, rs2

Carry-less multiply (high-part)

✓

✓

✓

clmulr rd, rs1, rs2

Carry-less multiply (reversed)

✓

✓

✓

clz rd, rs

Count leading zero bits

✓

✓

clzw rd, rs

Count leading zero bits in word

✓

✓

✓

cpop rd, rs

Count set bits

✓

✓

cpopw rd, rs

Count set bits in word

✓

✓

✓

ctz rd, rs

Count trailing zero bits

✓

✓

ctzw rd, rs

Count trailing zero bits in word

✓

✓

✓

max rd, rs1, rs2

Maximum

✓

✓

✓

maxu rd, rs1, rs2

Unsigned maximum

✓

✓

✓

min rd, rs1, rs2

Minimum

✓

✓

✓

minu rd, rs1, rs2

Unsigned minimum

✓

✓

✓

orc.b rd, rs

Bitwise OR-Combine, byte granule

✓

✓

✓

orn rd, rs1, rs2

OR with inverted operand

✓

✓

✓

rev8 rd, rs

Byte-reverse register

✓

✓

✓

rol rd, rs1, rs2

Rotate left (Register)

✓

✓

rolw rd, rs1, rs2

Rotate Left Word (Register)

✓

✓

✓

ror rd, rs1, rs2

Rotate right (Register)

✓

✓

✓

rori rd, rs1, shamt

Rotate right (Immediate)

✓

✓

roriw rd, rs1, shamt

Rotate right Word (Immediate)

✓

✓

rorw rd, rs1, rs2

Rotate right Word (Register)

✓

✓

✓

bclr rd, rs1, rs2

Single-Bit Clear (Register)

✓

✓

✓

bclri rd, rs1, imm

Single-Bit Clear (Immediate)

✓

✓

✓

bext rd, rs1, rs2

Single-Bit Extract (Register)

✓

✓

✓

bexti rd, rs1, imm

Single-Bit Extract (Immediate)

✓

✓

✓

binv rd, rs1, rs2

Single-Bit Invert (Register)

✓

✓

✓

binvi rd, rs1, imm

Single-Bit Invert (Immediate)

✓

✓

✓

bset rd, rs1, rs2

Single-Bit Set (Register)

✓

✓

✓

bseti rd, rs1, imm

Single-Bit Set (Immediate)

✓

✓

✓

sext.b rd, rs

Sign-extend byte

✓

✓

✓

sext.h rd, rs

Sign-extend halfword

✓

✓

✓

sh1add rd, rs1, rs2

Shift left by 1 and add

✓

✓

sh1add.uw rd, rs1, rs2

Shift unsigned word left by 1 and add

✓

✓

✓

sh2add rd, rs1, rs2

Shift left by 2 and add

✓

✓

sh2add.uw rd, rs1, rs2

Shift unsigned word left by 2 and add

✓

✓

✓

sh3add rd, rs1, rs2

Shift left by 3 and add

✓

✓

sh3add.uw rd, rs1, rs2

Shift unsigned word left by 3 and add

✓

✓

slli.uw rd, rs1, imm

Shift-left unsigned word (Immediate)

✓

✓

✓

xnor rd, rs1, rs2

Exclusive NOR

✓

✓

✓

zext.h rd, rs

Zero-extend halfword

✓

30.4.1. Zba: Address generation

The Zba instructions can be used to accelerate the generation of addresses that index into arrays of basic types (halfword, word, doubleword) using both unsigned word-sized and XLEN-sized indices: a shifted index is added to a base address.

The shift and add instructions do a left shift of 1, 2, or 3 because these are commonly found in real-world code and because they can be implemented with a minimal amount of additional hardware beyond that of the simple adder. This avoids lengthening the critical path in implementations.

While the shift and add instructions are limited to a maximum left shift of 3, the slli instruction (from the base ISA) can be used to perform similar shifts for indexing into arrays of wider elements. The slli.uw — added in this extension — can be used when the index is to be interpreted as an unsigned word.

The following instructions comprise the Zba extension:

RV32 RV64 Mnemonic Instruction

✓

add.uw rd, rs1, rs2

Add unsigned word

✓

✓

sh1add rd, rs1, rs2

Shift left by 1 and add

✓

sh1add.uw rd, rs1, rs2

Shift unsigned word left by 1 and add

✓

✓

sh2add rd, rs1, rs2

Shift left by 2 and add

✓

sh2add.uw rd, rs1, rs2

Shift unsigned word left by 2 and add

✓

✓

sh3add rd, rs1, rs2

Shift left by 3 and add

✓

sh3add.uw rd, rs1, rs2

Shift unsigned word left by 3 and add

✓

slli.uw rd, rs1, imm

Shift-left unsigned word (Immediate)

30.4.2. Zbb: Basic bit-manipulation

30.4.2.1. Logical with negate
RV32 RV64 Mnemonic Instruction

✓

✓

andn rd, rs1, rs2

AND with inverted operand

✓

✓

orn rd, rs1, rs2

OR with inverted operand

✓

✓

xnor rd, rs1, rs2

Exclusive NOR
Implementation Hint

The Logical with Negate instructions can be implemented by inverting the rs2 inputs to the base-required AND, OR, and XOR logic instructions. In some implementations, the inverter on rs2 used for subtraction can be reused for this purpose.
30.4.2.2. Count leading/trailing zero bits
RV32 RV64 Mnemonic Instruction

✓

✓

clz rd, rs

Count leading zero bits

✓

clzw rd, rs

Count leading zero bits in word

✓

✓

ctz rd, rs

Count trailing zero bits

✓

ctzw rd, rs

Count trailing zero bits in word
30.4.2.3. Count population

These instructions count the number of set bits (1-bits). This is also commonly referred to as population count.

RV32 RV64 Mnemonic Instruction

✓

✓

cpop rd, rs

Count set bits

✓

cpopw rd, rs

Count set bits in word
30.4.2.4. Integer minimum/maximum

The integer minimum/maximum instructions are arithmetic R-type instructions that return the smaller/larger of two operands.

RV32 RV64 Mnemonic Instruction

✓

✓

max rd, rs1, rs2

Maximum

✓

✓

maxu rd, rs1, rs2

Unsigned maximum

✓

✓

min rd, rs1, rs2

Minimum

✓

✓

minu rd, rs1, rs2

Unsigned minimum
30.4.2.5. Sign extension and zero extension

These instructions perform the sign extension or zero extension of the least significant 8 bits or 16 bits of the source register.

These instructions replace the generalized idioms slli rD,rS,(XLEN-<size>) + srli (for zero extension) or slli + srai (for sign extension) for the sign extension of 8-bit and 16-bit quantities, and for the zero extension of 16-bit quantities.

RV32 RV64 Mnemonic Instruction

✓

✓

sext.b rd, rs

Sign-extend byte

✓

✓

sext.h rd, rs

Sign-extend halfword

✓

✓

zext.h rd, rs

Zero-extend halfword
30.4.2.6. Bitwise rotation

Bitwise rotation instructions are similar to the shift-logical operations from the base spec. However, where the shift-logical instructions shift in zeros, the rotate instructions shift in the bits that were shifted out of the other side of the value. Such operations are also referred to as ‘circular shifts’.

RV32 RV64 Mnemonic Instruction

✓

✓

rol rd, rs1, rs2

Rotate left (Register)

✓

rolw rd, rs1, rs2

Rotate Left Word (Register)

✓

✓

ror rd, rs1, rs2

Rotate right (Register)

✓

✓

rori rd, rs1, shamt

Rotate right (Immediate)

✓

roriw rd, rs1, shamt

Rotate right Word (Immediate)

✓

rorw rd, rs1, rs2

Rotate right Word (Register)
Architecture Explanation

The rotate instructions were included to replace a common four-instruction sequence to achieve the same effect (neg; sll/srl; srl/sll; or)
30.4.2.7. OR Combine

orc.b sets the bits of each byte in the result rd to all zeros if no bit within the respective byte of rs is set, or to all ones if any bit within the respective byte of rs is set.

One use-case is string-processing functions, such as strlen and strcpy, which can use orc.b to test for the terminating zero byte by counting the set bits in leading non-zero bytes in a word.

RV32 RV64 Mnemonic Instruction

✓

✓

orc.b rd, rs

Bitwise OR-Combine, byte granule
30.4.2.8. Byte-reverse

rev8 reverses the byte-ordering of rs.

RV32 RV64 Mnemonic Instruction

✓

✓

rev8 rd, rs

Byte-reverse register

30.4.3. Zbc: Carry-less multiplication

Carry-less multiplication is the multiplication in the polynomial ring over GF(2).

clmul produces the lower half of the carry-less product and clmulh produces the upper half of the 2✕XLEN carry-less product.

clmulr produces bits 2✕XLEN−2:XLEN-1 of the 2✕XLEN carry-less product.

RV32 RV64 Mnemonic Instruction

✓

✓

clmul rd, rs1, rs2

Carry-less multiply (low-part)

✓

✓

clmulh rd, rs1, rs2

Carry-less multiply (high-part)

✓

✓

clmulr rd, rs1, rs2

Carry-less multiply (reversed)

30.4.4. Zbs: Single-bit instructions

The single-bit instructions provide a mechanism to set, clear, invert, or extract a single bit in a register. The bit is specified by its index.

RV32 RV64 Mnemonic Instruction

✓

✓

bclr rd, rs1, rs2

Single-Bit Clear (Register)

✓

✓

bclri rd, rs1, imm

Single-Bit Clear (Immediate)

✓

✓

bext rd, rs1, rs2

Single-Bit Extract (Register)

✓

✓

bexti rd, rs1, imm

Single-Bit Extract (Immediate)

✓

✓

binv rd, rs1, rs2

Single-Bit Invert (Register)

✓

✓

binvi rd, rs1, imm

Single-Bit Invert (Immediate)

✓

✓

bset rd, rs1, rs2

Single-Bit Set (Register)

✓

✓

bseti rd, rs1, imm

Single-Bit Set (Immediate)

30.4.5. Zbkb: Bit-manipulation for Cryptography

This extension contains instructions essential for implementing common operations in cryptographic workloads.

RV32 RV64 Mnemonic Instruction

✓

✓

rol

Rotate left (Register)

✓

rolw

Rotate Left Word (Register)

✓

✓

ror

Rotate right (Register)

✓

✓

rori

Rotate right (Immediate)

✓

roriw

Rotate right Word (Immediate)

✓

rorw

Rotate right Word (Register)

✓

✓

andn

AND with inverted operand

✓

✓

orn

OR with inverted operand

✓

✓

xnor

Exclusive NOR

✓

✓

pack

Pack low halves of registers

✓

✓

packh

Pack low bytes of registers

✓

packw

Pack low 16-bits of registers (RV64)

✓

✓

brev8

Reverse bits in bytes

✓

✓

rev8

Byte-reverse register

✓

zip

Bit interleave

✓

unzip

Bit deinterleave

30.4.6. Zbkc: Carry-less multiplication for Cryptography

Carry-less multiplication is the multiplication in the polynomial ring over GF(2). This is a critical operation in some cryptographic workloads, particularly the AES-GCM authenticated encryption scheme. This extension provides only the instructions needed to efficiently implement the GHASH operation, which is part of this workload.

RV32 RV64 Mnemonic Instruction

✓

✓

clmul rd, rs1, rs2

Carry-less multiply (low-part)

✓

✓

clmulh rd, rs1, rs2

Carry-less multiply (high-part)

30.4.7. Zbkx: Crossbar permutations

These instructions implement a "lookup table" for 4 and 8 bit elements inside the general purpose registers. rs1 is used as a vector of N-bit words, and rs2 as a vector of N-bit indices into rs1. Elements in rs1 are replaced by the indexed element in rs2, or zero if the index into rs2 is out of bounds.

These instructions are useful for expressing N-bit to N-bit boolean operations, and implementing cryptographic code with secret dependent memory accesses (particularly SBoxes) such that the execution latency does not depend on the (secret) data being operated on.

RV32 RV64 Mnemonic Instruction

✓

✓

xperm4 rd, rs1, rs2

Crossbar permutation (nibbles)

✓

✓

xperm8 rd, rs1, rs2

Crossbar permutation (bytes)

30.5. Instructions (in alphabetical order)

30.5.1. add.uw

Synopsis

Add unsigned word

Mnemonic

add.uw rd, rs1, rs2

Pseudoinstructions

zext.w rd, rs1 → add.uw rd, rs1, zero

Encoding
Diagram
Description

This instruction performs an XLEN-wide addition between rs2 and the zero-extended least-significant word of rs1.

Operation
let base = X(rs2);
let index = EXTZ(X(rs1)[31..0]);

X(rd) = base + index;
Included in
Extension Minimum version Lifecycle state

Zba (Address generation instructions)

0.93

Ratified

30.5.2. andn

Synopsis

AND with inverted operand

Mnemonic

andn rd, rs1, rs2

Encoding
Diagram
Description

This instruction performs the bitwise logical AND operation between rs1 and the bitwise inversion of rs2.

Operation
X(rd) = X(rs1) & ~X(rs2);
Included in
Extension Minimum version Lifecycle state

Zbb (Basic bit-manipulation)

v1.0

Ratified

Zbkb (Bit-manipulation for Cryptography)

v1.0

Ratified

30.5.3. bclr

Synopsis

Single-Bit Clear (Register)

Mnemonic

bclr rd, rs1, rs2

Encoding
Diagram
Description

This instruction returns rs1 with a single bit cleared at the index specified in rs2. The index is read from the lower log2(XLEN) bits of rs2.

Operation
let index = X(rs2) & (XLEN - 1);
X(rd) = X(rs1) & ~(1 << index)
Included in
Extension Minimum version Lifecycle state

Zbs (Single-bit instructions)

v1.0

Ratified

30.5.4. bclri

Synopsis

Single-Bit Clear (Immediate)

Mnemonic

bclri rd, rs1, shamt

Encoding (RV32)
Diagram
Encoding (RV64)
Diagram
Description

This instruction returns rs1 with a single bit cleared at the index specified in shamt. The index is read from the lower log2(XLEN) bits of shamt. For RV32, the encodings corresponding to shamt[5]=1 are reserved.

Operation
let index = shamt & (XLEN - 1);
X(rd) = X(rs1) & ~(1 << index)
Included in
Extension Minimum version Lifecycle state

Zbs (Single-bit instructions)

v1.0

Ratified

30.5.5. bext

Synopsis

Single-Bit Extract (Register)

Mnemonic

bext rd, rs1, rs2

Encoding
Diagram
Description

This instruction returns a single bit extracted from rs1 at the index specified in rs2. The index is read from the lower log2(XLEN) bits of rs2.

Operation
let index = X(rs2) & (XLEN - 1);
X(rd) = (X(rs1) >> index) & 1;
Included in
Extension Minimum version Lifecycle state

Zbs (Single-bit instructions)

v1.0

Ratified

30.5.6. bexti

Synopsis

Single-Bit Extract (Immediate)

Mnemonic

bexti rd, rs1, shamt

Encoding (RV32)
Diagram
Encoding (RV64)
Diagram
Description

This instruction returns a single bit extracted from rs1 at the index specified in rs2. The index is read from the lower log2(XLEN) bits of shamt. For RV32, the encodings corresponding to shamt[5]=1 are reserved.

Operation
let index = shamt & (XLEN - 1);
X(rd) = (X(rs1) >> index) & 1;
Included in
Extension Minimum version Lifecycle state

Zbs (Single-bit instructions)

v1.0

Ratified

30.5.7. binv

Synopsis

Single-Bit Invert (Register)

Mnemonic

binv rd, rs1, rs2

Encoding
Diagram
Description

This instruction returns rs1 with a single bit inverted at the index specified in rs2. The index is read from the lower log2(XLEN) bits of rs2.

Operation
let index = X(rs2) & (XLEN - 1);
X(rd) = X(rs1) ^ (1 << index)
Included in
Extension Minimum version Lifecycle state

Zbs (Single-bit instructions)

v1.0

Ratified

30.5.8. binvi

Synopsis

Single-Bit Invert (Immediate)

Mnemonic

binvi rd, rs1, shamt

Encoding (RV32)
Diagram
Encoding (RV64)
Diagram
Description

This instruction returns rs1 with a single bit inverted at the index specified in shamt. The index is read from the lower log2(XLEN) bits of shamt. For RV32, the encodings corresponding to shamt[5]=1 are reserved.

Operation
let index = shamt & (XLEN - 1);
X(rd) = X(rs1) ^ (1 << index)
Included in
Extension Minimum version Lifecycle state

Zbs (Single-bit instructions)

v1.0

Ratified

30.5.9. bset

Synopsis

Single-Bit Set (Register)

Mnemonic

bset rd, rs1,rs2

Encoding
Diagram
Description

This instruction returns rs1 with a single bit set at the index specified in rs2. The index is read from the lower log2(XLEN) bits of rs2.

Operation
let index = X(rs2) & (XLEN - 1);
X(rd) = X(rs1) | (1 << index)
Included in
Extension Minimum version Lifecycle state

Zbs (Single-bit instructions)

v1.0

Ratified

30.5.10. bseti

Synopsis

Single-Bit Set (Immediate)

Mnemonic

bseti rd, rs1,shamt

Encoding (RV32)
Diagram
Encoding (RV64)
Diagram
Description

This instruction returns rs1 with a single bit set at the index specified in shamt. The index is read from the lower log2(XLEN) bits of shamt. For RV32, the encodings corresponding to shamt[5]=1 are reserved.

Operation
let index = shamt & (XLEN - 1);
X(rd) = X(rs1) | (1 << index)
Included in
Extension Minimum version Lifecycle state

Zbs (Single-bit instructions)

v1.0

Ratified

30.5.11. clmul

Synopsis

Carry-less multiply (low-part)

Mnemonic

clmul rd, rs1, rs2

Encoding
Diagram
Description

clmul produces the lower half of the 2·XLEN carry-less product.

Operation
let rs1_val = X(rs1);
let rs2_val = X(rs2);
let output : xlenbits = 0;

foreach (i from 0 to (xlen - 1) by 1) {
   output = if   ((rs2_val >> i) & 1)
            then output ^ (rs1_val << i);
            else output;
}

X[rd] = output
Included in
Extension Minimum version Lifecycle state

Zbc (Carry-less multiplication)

v1.0

Ratified

Zbkc (Carry-less multiplication for Cryptography)

v1.0

Ratified

30.5.12. clmulh

Synopsis

Carry-less multiply (high-part)

Mnemonic

clmulh rd, rs1, rs2

Encoding
Diagram
Description

clmulh produces the upper half of the 2·XLEN carry-less product.

Operation
let rs1_val = X(rs1);
let rs2_val = X(rs2);
let output : xlenbits = 0;

foreach (i from 1 to xlen by 1) {
   output = if   ((rs2_val >> i) & 1)
            then output ^ (rs1_val >> (xlen - i));
            else output;
}

X[rd] = output
Included in
Extension Minimum version Lifecycle state

Zbc (Carry-less multiplication)

v1.0

Ratified

Zbkc (Carry-less multiplication for Cryptography)

v1.0

Ratified

30.5.13. clmulr

Synopsis

Carry-less multiply (reversed)

Mnemonic

clmulr rd, rs1, rs2

Encoding
Diagram
Description

clmulr produces bits 2·XLEN−2:XLEN-1 of the 2·XLEN carry-less product.

Operation
let rs1_val = X(rs1);
let rs2_val = X(rs2);
let output : xlenbits = 0;

foreach (i from 0 to (xlen - 1) by 1) {
   output = if   ((rs2_val >> i) & 1)
            then output ^ (rs1_val >> (xlen - i - 1));
            else output;
}

X[rd] = output
Note

The clmulr instruction is used to accelerate CRC calculations. The r in the instruction’s mnemonic stands for reversed, as the instruction is equivalent to bit-reversing the inputs, performing a clmul, then bit-reversing the output.
Included in
Extension Minimum version Lifecycle state

Zbc (Carry-less multiplication)

v1.0

Ratified

30.5.14. clz

Synopsis

Count leading zero bits

Mnemonic

clz rd, rs

Encoding
Diagram
Description

This instruction counts the number of 0’s before the first 1, starting at the most-significant bit (i.e., XLEN-1) and progressing to bit 0. Accordingly, if the input is 0, the output is XLEN, and if the most-significant bit of the input is a 1, the output is 0.

Operation
val HighestSetBit : forall ('N : Int), 'N >= 0. bits('N) -> int

function HighestSetBit x = {
  foreach (i from (xlen - 1) to 0 by 1 in dec)
    if [x[i]] == 0b1 then return(i) else ();
  return -1;
}

let rs = X(rs);
X[rd] = (xlen - 1) - HighestSetBit(rs);
Included in
Extension Minimum version Lifecycle state

Zbb (Basic bit-manipulation)

v1.0

Ratified

30.5.15. clzw

Synopsis

Count leading zero bits in word

Mnemonic

clzw rd, rs

Encoding
Diagram
Description

This instruction counts the number of 0’s before the first 1 starting at bit 31 and progressing to bit 0. Accordingly, if the least-significant word is 0, the output is 32, and if the most-significant bit of the word (i.e., bit 31) is a 1, the output is 0.

Operation
val HighestSetBit32 : forall ('N : Int), 'N >= 0. bits('N) -> int

function HighestSetBit32 x = {
  foreach (i from 31 to 0 by 1 in dec)
    if [x[i]] == 0b1 then return(i) else ();
  return -1;
}

let rs = X(rs);
X[rd] = 31 - HighestSetBit(rs);
Included in
Extension Minimum version Lifecycle state

Zbb (Basic bit-manipulation)

v1.0

Ratified

30.5.16. cpop

Synopsis

Count set bits

Mnemonic

cpop rd, rs

Encoding
Diagram
Description

This instructions counts the number of 1’s (i.e., set bits) in the source register.

Operation
let bitcount = 0;
let rs = X(rs);

foreach (i from 0 to (xlen - 1) in inc)
    if rs[i] == 0b1 then bitcount = bitcount + 1 else ();

X[rd] = bitcount
Software Hint

This operations is known as population count, popcount, sideways sum, bit summation, or Hamming weight.

The GCC builtin function __builtin_popcount (unsigned int x) is implemented by cpop on RV32 and by cpopw on RV64. The GCC builtin function __builtin_popcountl (unsigned long x) for LP64 is implemented by cpop on RV64.
Included in
Extension Minimum version Lifecycle state

Zbb (Basic bit-manipulation)

v1.0

Ratified

30.5.17. cpopw

Synopsis

Count set bits in word

Mnemonic

cpopw rd, rs

Encoding
Diagram
Description

This instructions counts the number of 1’s (i.e., set bits) in the least-significant word of the source register.

Operation
let bitcount = 0;
let val = X(rs);

foreach (i from 0 to 31 in inc)
    if val[i] == 0b1 then bitcount = bitcount + 1 else ();

X[rd] = bitcount
Included in
Extension Minimum version Lifecycle state

Zbb (Basic bit-manipulation)

v1.0

Ratified

30.5.18. ctz

Synopsis

Count trailing zeros

Mnemonic

ctz rd, rs

Encoding
Diagram
Description

This instruction counts the number of 0’s before the first 1, starting at the least-significant bit (i.e., 0) and progressing to the most-significant bit (i.e., XLEN-1). Accordingly, if the input is 0, the output is XLEN, and if the least-significant bit of the input is a 1, the output is 0.

Operation
val LowestSetBit : forall ('N : Int), 'N >= 0. bits('N) -> int

function LowestSetBit x = {
  foreach (i from 0 to (xlen - 1) by 1 in dec)
    if [x[i]] == 0b1 then return(i) else ();
  return xlen;
}

let rs = X(rs);
X[rd] = LowestSetBit(rs);
Included in
Extension Minimum version Lifecycle state

Zbb (Basic bit-manipulation)

v1.0

Ratified

30.5.19. ctzw

Synopsis

Count trailing zero bits in word

Mnemonic

ctzw rd, rs

Encoding
Diagram
Description

This instruction counts the number of 0’s before the first 1, starting at the least-significant bit (i.e., 0) and progressing to the most-significant bit of the least-significant word (i.e., 31). Accordingly, if the least-significant word is 0, the output is 32, and if the least-significant bit of the input is a 1, the output is 0.

Operation
val LowestSetBit32 : forall ('N : Int), 'N >= 0. bits('N) -> int

function LowestSetBit32 x = {
  foreach (i from 0 to 31 by 1 in dec)
    if [x[i]] == 0b1 then return(i) else ();
  return 32;
}

let rs = X(rs);
X[rd] = LowestSetBit32(rs);
Included in
Extension Minimum version Lifecycle state

Zbb (Basic bit-manipulation)

v1.0

Ratified

30.5.20. max

Synopsis

Maximum

Mnemonic

max rd, rs1, rs2

Encoding
Diagram
Description

This instruction returns the larger of two signed integers.

Operation
let rs1_val = X(rs1);
let rs2_val = X(rs2);

let result = if   rs1_val <_s rs2_val
             then rs2_val
             else rs1_val;

X(rd) = result;
Software Hint

Calculating the absolute value of a signed integer can be performed using the following sequence: neg rD,rS followed by max rD,rS,rD. When using this common sequence, it is suggested that they are scheduled with no intervening instructions so that implementations that are so optimized can fuse them together.
Included in
Extension Minimum version Lifecycle state

Zbb (Basic bit-manipulation)

v1.0

Ratified

30.5.21. maxu

Synopsis

Unsigned maximum

Mnemonic

maxu rd, rs1, rs2

Encoding
Diagram
Description

This instruction returns the larger of two unsigned integers.

Operation
let rs1_val = X(rs1);
let rs2_val = X(rs2);

let result = if   rs1_val <_u rs2_val
             then rs2_val
             else rs1_val;

X(rd) = result;
Included in
Extension Minimum version Lifecycle state

Zbb (Basic bit-manipulation)

v1.0

Ratified

30.5.22. min

Synopsis

Minimum

Mnemonic

min rd, rs1, rs2

Encoding
Diagram
Description

This instruction returns the smaller of two signed integers.

Operation
let rs1_val = X(rs1);
let rs2_val = X(rs2);

let result = if   rs1_val <_s rs2_val
             then rs1_val
             else rs2_val;

X(rd) = result;
Included in
Extension Minimum version Lifecycle state

Zbb (Basic bit-manipulation)

v1.0

Ratified

30.5.23. minu

Synopsis

Unsigned minimum

Mnemonic

minu rd, rs1, rs2

Encoding
Diagram
Description

This instruction returns the smaller of two unsigned integers.

Operation
let rs1_val = X(rs1);
let rs2_val = X(rs2);

let result = if   rs1_val <_u rs2_val
             then rs1_val
             else rs2_val;

X(rd) = result;
Included in
Extension Minimum version Lifecycle state

Zbb (Basic bit-manipulation)

v1.0

Ratified

30.5.24. orc.b

Synopsis

Bitwise OR-Combine, byte granule

Mnemonic

orc.b rd, rs

Encoding
Diagram
Description

Combines the bits within each byte using bitwise logical OR. This sets the bits of each byte in the result rd to all zeros if no bit within the respective byte of rs is set, or to all ones if any bit within the respective byte of rs is set.

Operation
let input = X(rs);
let output : xlenbits = 0;

foreach (i from 0 to (xlen - 8) by 8) {
   output[(i + 7)..i] = if   input[(i + 7)..i] == 0
                        then 0b00000000
                        else 0b11111111;
}

X[rd] = output;
Included in
Extension Minimum version Lifecycle state

Zbb (Basic bit-manipulation)

v1.0

Ratified

30.5.25. orn

Synopsis

OR with inverted operand

Mnemonic

orn rd, rs1, rs2

Encoding
Diagram
Description

This instruction performs the bitwise logical OR operation between rs1 and the bitwise inversion of rs2.

Operation
X(rd) = X(rs1) | ~X(rs2);
Included in
Extension Minimum version Lifecycle state

Zbb (Basic bit-manipulation)

v1.0

Ratified

Zbkb (Bit-manipulation for Cryptography)

v1.0

Ratified

30.5.26. pack

Synopsis

Pack the low halves of rs1 and rs2 into rd.

Mnemonic

pack rd, rs1, rs2

Encoding
Diagram
Description

The pack instruction packs the XLEN/2-bit lower halves of rs1 and rs2 into rd, with rs1 in the lower half and rs2 in the upper half.

Operation
let lo_half : bits(xlen/2) = X(rs1)[xlen/2-1..0];
let hi_half : bits(xlen/2) = X(rs2)[xlen/2-1..0];
X(rd) = EXTZ(hi_half @ lo_half);
Included in
Extension Minimum version Lifecycle state

Zbkb (Bit-manipulation for Cryptography)

v1.0

Ratified
For RV32, the pack instruction with rs2=x0 is the zext.h instruction. Hence, for RV32, any extension that contains the pack instruction also contains the zext.h instruction (but not necessarily the c.zext.h instruction, which is only guaranteed to exist if both the Zcb and Zbb extensions are implemented).

30.5.27. packh

Synopsis

Pack the low bytes of rs1 and rs2 into rd.

Mnemonic

packh rd, rs1, rs2

Encoding
Diagram
Description

And the packh instruction packs the least-significant bytes of rs1 and rs2 into the 16 least-significant bits of rd, zero extending the rest of rd.

Operation
let lo_half : bits(8) = X(rs1)[7..0];
let hi_half : bits(8) = X(rs2)[7..0];
X(rd) = EXTZ(hi_half @ lo_half);
Included in
Extension Minimum version Lifecycle state

Zbkb (Bit-manipulation for Cryptography)

v1.0

Ratified

30.5.28. packw

Synopsis

Pack the low 16-bits of rs1 and rs2 into rd on RV64.

Mnemonic

packw rd, rs1, rs2

Encoding
Diagram
Description

This instruction packs the low 16 bits of rs1 and rs2 into the 32 least-significant bits of rd, sign extending the 32-bit result to the rest of rd. This instruction only exists on RV64 based systems.

Operation
let lo_half : bits(16) = X(rs1)[15..0];
let hi_half : bits(16) = X(rs2)[15..0];
X(rd) = EXTS(hi_half @ lo_half);
Included in
Extension Minimum version Lifecycle state

Zbkb (Bit-manipulation for Cryptography)

v1.0

Ratified
For RV64, the packw instruction with rs2=x0 is the zext.h instruction. Hence, for RV64, any extension that contains the packw instruction also contains the zext.h instruction (but not necessarily the c.zext.h instruction, which is only guaranteed to exist if both the Zcb and Zbb extensions are implemented).

30.5.29. rev8

Synopsis

Byte-reverse register

Mnemonic

rev8 rd, rs

Encoding (RV32)
Diagram
Encoding (RV64)
Diagram
Description

This instruction reverses the order of the bytes in rs.

Operation
let input = X(rs);
let output : xlenbits = 0;
let j = xlen - 1;

foreach (i from 0 to (xlen - 8) by 8) {
   output[i..(i + 7)] = input[(j - 7)..j];
   j = j - 8;
}

X[rd] = output
Note

The rev8 mnemonic corresponds to different instruction encodings in RV32 and RV64.
Software Hint

The byte-reverse operation is only available for the full register width. To emulate word-sized and halfword-sized byte-reversal, perform a rev8 rd,rs followed by a srai rd,rd,K, where K is XLEN-32 and XLEN-16, respectively.
Included in
Extension Minimum version Lifecycle state

Zbb (Basic bit-manipulation)

v1.0

Ratified

Zbkb (Bit-manipulation for Cryptography)

v1.0

Ratified

30.5.30. brev8

Synopsis

Reverse the bits in each byte of a source register.

Mnemonic

brev8 rd, rs

Encoding
Diagram
Description

This instruction reverses the order of the bits in every byte of a register.

Operation
result : xlenbits = EXTZ(0b0);
foreach (i from 0 to sizeof(xlen) by 8) {
    result[i+7..i] = reverse_bits_in_byte(X(rs1)[i+7..i]);
};
X(rd) = result;
Included in
Extension Minimum version Lifecycle state

Zbkb (Bit-manipulation for Cryptography)

v1.0

Ratified

30.5.31. rol

Synopsis

Rotate Left (Register)

Mnemonic

rol rd, rs1, rs2

Encoding
Diagram
Description

This instruction performs a rotate left of rs1 by the amount in least-significant log2(XLEN) bits of rs2.

Operation
let shamt = if   xlen == 32
            then X(rs2)[4..0]
            else X(rs2)[5..0];
let result = (X(rs1) << shamt) | (X(rs1) >> (xlen - shamt));

X(rd) = result;
Included in
Extension Minimum version Lifecycle state

Zbb (Basic bit-manipulation)

0.93

Ratified

Zbkb (Bit-manipulation for Cryptography)

v1.0

Ratified

30.5.32. rolw

Synopsis

Rotate Left Word (Register)

Mnemonic

rolw rd, rs1, rs2

Encoding
Diagram
Description

This instruction performs a rotate left on the least-significant word of rs1 by the amount in least-significant 5 bits of rs2. The resulting word value is sign-extended by copying bit 31 to all of the more-significant bits.

Operation
let rs1 = EXTZ(X(rs1)[31..0])
let shamt = X(rs2)[4..0];
let result = (rs1 << shamt) | (rs1 >> (32 - shamt));
X(rd) = EXTS(result[31..0]);
Included in
Extension Minimum version Lifecycle state

Zbb (Basic bit-manipulation)

0.93

Ratified

Zbkb (Bit-manipulation for Cryptography)

v1.0

Ratified

30.5.33. ror

Synopsis

Rotate Right

Mnemonic

ror rd, rs1, rs2

Encoding
Diagram
Description

This instruction performs a rotate right of rs1 by the amount in least-significant log2(XLEN) bits of rs2.

Operation
let shamt = if   xlen == 32
            then X(rs2)[4..0]
            else X(rs2)[5..0];
let result = (X(rs1) >> shamt) | (X(rs1) << (xlen - shamt));

X(rd) = result;
Included in
Extension Minimum version Lifecycle state

Zbb (Basic bit-manipulation)

0.93

Ratified

Zbkb (Bit-manipulation for Cryptography)

v1.0

Ratified

30.5.34. rori

Synopsis

Rotate Right (Immediate)

Mnemonic

rori rd, rs1, shamt

Encoding (RV32)
Diagram
Encoding (RV64)
Diagram
Description

This instruction performs a rotate right of rs1 by the amount in the least-significant log2(XLEN) bits of shamt. For RV32, the encodings corresponding to shamt[5]=1 are reserved.

Operation
let shamt = if   xlen == 32
            then shamt[4..0]
            else shamt[5..0];
let result = (X(rs1) >> shamt) | (X(rs1) << (xlen - shamt));

X(rd) = result;
Included in
Extension Minimum version Lifecycle state

Zbb (Basic bit-manipulation)

0.93

Ratified

Zbkb (Bit-manipulation for Cryptography)

v1.0

Ratified

30.5.35. roriw

Synopsis

Rotate Right Word by Immediate

Mnemonic

roriw rd, rs1, shamt

Encoding
Diagram
Description

This instruction performs a rotate right on the least-significant word of rs1 by the amount in the least-significant log2(XLEN) bits of shamt. The resulting word value is sign-extended by copying bit 31 to all of the more-significant bits.

Operation
let rs1_data = EXTZ(X(rs1)[31..0];
let result = (rs1_data >> shamt) | (rs1_data << (32 - shamt));
X(rd) = EXTS(result[31..0]);
Included in
Extension Minimum version Lifecycle state

Zbb (Basic bit-manipulation)

0.93

Ratified

Zbkb (Bit-manipulation for Cryptography)

v1.0

Ratified

30.5.36. rorw

Synopsis

Rotate Right Word (Register)

Mnemonic

rorw rd, rs1, rs2

Encoding
Diagram
Description

This instruction performs a rotate right on the least-significant word of rs1 by the amount in least-significant 5 bits of rs2. The resultant word is sign-extended by copying bit 31 to all of the more-significant bits.

Operation
let rs1 = EXTZ(X(rs1)[31..0])
let shamt = X(rs2)[4..0];
let result = (rs1 >> shamt) | (rs1 << (32 - shamt));
X(rd) = EXTS(result);
Included in
Extension Minimum version Lifecycle state

Zbb (Basic bit-manipulation)

0.93

Ratified

Zbkb (Bit-manipulation for Cryptography)

v1.0

Ratified

30.5.37. sext.b

Synopsis

Sign-extend byte

Mnemonic

sext.b rd, rs

Encoding
Diagram
Description

This instruction sign-extends the least-significant byte in the source to XLEN by copying the most-significant bit in the byte (i.e., bit 7) to all of the more-significant bits.

Operation
X(rd) = EXTS(X(rs)[7..0]);
Included in
Extension Minimum version Lifecycle state

Zbb (Basic bit-manipulation)

0.93

Ratified

30.5.38. sext.h

Synopsis

Sign-extend halfword

Mnemonic

sext.h rd, rs

Encoding
Diagram
Description

This instruction sign-extends the least-significant halfword in rs to XLEN by copying the most-significant bit in the halfword (i.e., bit 15) to all of the more-significant bits.

Operation
X(rd) = EXTS(X(rs)[15..0]);
Included in
Extension Minimum version Lifecycle state

Zbb (Basic bit-manipulation)

0.93

Ratified

30.5.39. sh1add

Synopsis

Shift left by 1 and add

Mnemonic

sh1add rd, rs1, rs2

Encoding
Diagram
Description

This instruction shifts rs1 to the left by 1 bit and adds it to rs2.

Operation
X(rd) = X(rs2) + (X(rs1) << 1);
Included in
Extension Minimum version Lifecycle state

Zba (Address generation instructions)

0.93

Ratified

30.5.40. sh1add.uw

Synopsis

Shift unsigned word left by 1 and add

Mnemonic

sh1add.uw rd, rs1, rs2

Encoding
Diagram
Description

This instruction performs an XLEN-wide addition of two addends. The first addend is rs2. The second addend is the unsigned value formed by extracting the least-significant word of rs1 and shifting it left by 1 place.

Operation
let base = X(rs2);
let index = EXTZ(X(rs1)[31..0]);

X(rd) = base + (index << 1);
Included in
Extension Minimum version Lifecycle state

Zba (Address generation instructions)

0.93

Ratified

30.5.41. sh2add

Synopsis

Shift left by 2 and add

Mnemonic

sh2add rd, rs1, rs2

Encoding
Diagram
Description

This instruction shifts rs1 to the left by 2 places and adds it to rs2.

Operation
X(rd) = X(rs2) + (X(rs1) << 2);
Included in
Extension Minimum version Lifecycle state

Zba (Address generation instructions)

0.93

Ratified

30.5.42. sh2add.uw

Synopsis

Shift unsigned word left by 2 and add

Mnemonic

sh2add.uw rd, rs1, rs2

Encoding
Diagram
Description

This instruction performs an XLEN-wide addition of two addends. The first addend is rs2. The second addend is the unsigned value formed by extracting the least-significant word of rs1 and shifting it left by 2 places.

Operation
let base = X(rs2);
let index = EXTZ(X(rs1)[31..0]);

X(rd) = base + (index << 2);
Included in
Extension Minimum version Lifecycle state

Zba (Address generation instructions)

0.93

Ratified

30.5.43. sh3add

Synopsis

Shift left by 3 and add

Mnemonic

sh3add rd, rs1, rs2

Encoding
Diagram
Description

This instruction shifts rs1 to the left by 3 places and adds it to rs2.

Operation
X(rd) = X(rs2) + (X(rs1) << 3);
Included in
Extension Minimum version Lifecycle state

Zba (Address generation instructions)

0.93

Ratified

30.5.44. sh3add.uw

Synopsis

Shift unsigned word left by 3 and add

Mnemonic

sh3add.uw rd, rs1, rs2

Encoding
Diagram
Description

This instruction performs an XLEN-wide addition of two addends. The first addend is rs2. The second addend is the unsigned value formed by extracting the least-significant word of rs1 and shifting it left by 3 places.

Operation
let base = X(rs2);
let index = EXTZ(X(rs1)[31..0]);

X(rd) = base + (index << 3);
Included in
Extension Minimum version Lifecycle state

Zba (Address generation instructions)

0.93

Ratified

30.5.45. slli.uw

Synopsis

Shift-left unsigned word (Immediate)

Mnemonic

slli.uw rd, rs1, shamt

Encoding
Diagram
Description

This instruction takes the least-significant word of rs1, zero-extends it, and shifts it left by the immediate.

Operation
X(rd) = (EXTZ(X(rs)[31..0]) << shamt);
Included in
Extension Minimum version Lifecycle state

Zba (Address generation instructions)

0.93

Ratified
Architecture Explanation

This instruction is the same as slli with zext.w performed on rs1 before shifting.

30.5.46. unzip

Synopsis

Place odd and even bits of the source register into upper and lower halves of the destination register, respectively.

Mnemonic

unzip rd, rs

Encoding
Diagram
Description

This instruction scatters all of the odd and even bits of a source word into the high and low halves of a destination word. It is the inverse of the zip instruction. This instruction is available only on RV32.

Operation
foreach (i from 0 to xlen/2-1) {
  X(rd)[i] = X(rs1)[2*i]
  X(rd)[i+xlen/2] = X(rs1)[2*i+1]
}
Software Hint

This instruction is useful for implementing the SHA3 cryptographic hash function on a 32-bit architecture, as it implements the bit-interleaving operation used to speed up the 64-bit rotations directly.
Included in
Extension Minimum version Lifecycle state

Zbkb (Bit-manipulation for Cryptography) (RV32)

v1.0

Ratified

30.5.47. xnor

Synopsis

Exclusive NOR

Mnemonic

xnor rd, rs1, rs2

Encoding
Diagram
Description

This instruction performs the bit-wise exclusive-NOR operation on rs1 and rs2.

Operation
X(rd) = ~(X(rs1) ^ X(rs2));
Included in
Extension Minimum version Lifecycle state

Zbb (Basic bit-manipulation)

0.93

Ratified

Zbkb (Bit-manipulation for Cryptography)

v1.0

Ratified

30.5.48. xperm8

Synopsis

Byte-wise lookup of indices into a vector in registers.

Mnemonic

xperm8 rd, rs1, rs2

Encoding
Diagram
Description

The xperm8 instruction operates on bytes. The rs1 register contains a vector of XLEN/8 8-bit elements. The rs2 register contains a vector of XLEN/8 8-bit indexes. The result is each element in rs2 replaced by the indexed element in rs1, or zero if the index into rs2 is out of bounds.

Operation
val xperm8_lookup : (bits(8), xlenbits) -> bits(8)
function xperm8_lookup (idx, lut) = {
    (lut >> (idx @ 0b000))[7..0]
}

function clause execute ( XPERM8 (rs2,rs1,rd)) = {
    result : xlenbits = EXTZ(0b0);
    foreach(i from 0 to xlen by 8) {
        result[i+7..i] = xperm8_lookup(X(rs2)[i+7..i], X(rs1));
    };
    X(rd) = result;
    RETIRE_SUCCESS
}
Included in
Extension Minimum version Lifecycle state

Zbkx (Crossbar permutations)

v1.0

Ratified

30.5.49. xperm4

Synopsis

Nibble-wise lookup of indices into a vector.

Mnemonic

xperm4 rd, rs1, rs2

Encoding
Diagram
Description

The xperm4 instruction operates on nibbles. The rs1 register contains a vector of XLEN/4 4-bit elements. The rs2 register contains a vector of XLEN/4 4-bit indexes. The result is each element in rs2 replaced by the indexed element in rs1, or zero if the index into rs2 is out of bounds.

Operation
val xperm4_lookup : (bits(4), xlenbits) -> bits(4)
function xperm4_lookup (idx, lut) = {
    (lut >> (idx @ 0b00))[3..0]
}

function clause execute ( XPERM4 (rs2,rs1,rd)) = {
    result : xlenbits = EXTZ(0b0);
    foreach(i from 0 to xlen by 4) {
        result[i+3..i] = xperm4_lookup(X(rs2)[i+3..i], X(rs1));
    };
    X(rd) = result;
    RETIRE_SUCCESS
}
Included in
Extension Minimum version Lifecycle state

Zbkx (Crossbar permutations)

v1.0

Ratified

30.5.50. zext.h

Synopsis

Zero-extend halfword

Mnemonic

zext.h rd, rs

Encoding (RV32)
Diagram
Encoding (RV64)
Diagram
Description

This instruction zero-extends the least-significant halfword of the source to XLEN by inserting 0’s into all of the bits more significant than 15.

Operation
X(rd) = EXTZ(X(rs)[15..0]);
Note

The zext.h mnemonic corresponds to different instruction encodings in RV32 and RV64.
Included in
Extension Minimum version Lifecycle state

Zbb (Basic bit-manipulation)

0.93

Ratified

30.5.51. zip

Synopsis

Interleave upper and lower halves of the source register into odd and even bits of the destination register, respectivley.

Mnemonic

zip rd, rs

Encoding
Diagram
Description

This instruction gathers bits from the high and low halves of the source word into odd/even bit positions in the destination word. It is the inverse of the unzip instruction. This instruction is available only on RV32.

Operation
foreach (i from 0 to xlen/2-1) {
  X(rd)[2*i] = X(rs1)[i]
  X(rd)[2*i+1] = X(rs1)[i+xlen/2]
}
Software Hint

This instruction is useful for implementing the SHA3 cryptographic hash function on a 32-bit architecture, as it implements the bit-interleaving operation used to speed up the 64-bit rotations directly.
Included in
Extension Minimum version Lifecycle state

Zbkb (Bit-manipulation for Cryptography) (RV32)

v1.0

Ratified

30.6. Software optimization guide

30.6.1. strlen

The orc.b instruction allows for the efficient detection of NUL bytes in an XLEN-sized chunk of data:

  • the result of orc.b on a chunk that does not contain any NUL bytes will be all-ones, and

  • after a bitwise-negation of the result of orc.b, the number of data bytes before the first NUL byte (if any) can be detected by ctz/clz (depending on the endianness of data).

A full example of a strlen function, which uses these techniques and also demonstrates the use of it for unaligned/partial data, is the following:

#include <sys/asm.h>

	.text
	.globl strlen
	.type  strlen, @function
strlen:
	andi	a3, a0, (SZREG-1)   // offset
	andi    a1, a0, -SZREG      // align pointer
.Lprologue:
	li      a4, SZREG
	sub     a4, a4, a3          // XLEN - offset
	slli	a3, a3, 3           // offset * 8
	REG_L   a2, 0(a1)           // chunk
	/*
	 * Shift the partial/unaligned chunk we loaded to remove the bytes
	 * from before the start of the string, adding NUL bytes at the end.
	 */
#if __BYTE_ORDER__ == __ORDER_LITTLE_ENDIAN__
	srl	a2, a2 ,a3          // chunk >> (offset * 8)
#else
	sll     a2, a2, a3
#endif
	orc.b   a2, a2
	not	a2, a2
	/*
	 * Non-NUL bytes in the string have been expanded to 0x00, while
 	 * NUL bytes have become 0xff.  Search for the first set bit
	 * (corresponding to a NUL byte in the original chunk).
	 */
#if __BYTE_ORDER__ == __ORDER_LITTLE_ENDIAN__
	ctz     a2, a2
#else
	clz     a2, a2
#endif
	/*
	 * The first chunk is special: compare against the number of valid
	 * bytes in this chunk.
	 */
	srli    a0, a2, 3
	bgtu    a4, a0, .Ldone
	addi    a3, a1, SZREG
	li      a4, -1
	.align 2
	/*
	 * Our critical loop is 4 instructions and processes data in 4 byte
	 * or 8 byte chunks.
	 */
.Lloop:
	REG_L   a2, SZREG(a1)
	addi    a1, a1, SZREG
	orc.b   a2, a2
	beq     a2, a4, .Lloop

.Lepilogue:
	not     a2, a2
#if __BYTE_ORDER__ == __ORDER_LITTLE_ENDIAN__
	ctz     a2, a2
#else
	clz     a2, a2
#endif
	sub     a1, a1, a3
	add	a0, a0, a1
	srli    a2, a2, 3
	add 	a0, a0, a2
.Ldone:
	ret

30.6.2. strcmp

#include <sys/asm.h>

  .text
  .globl strcmp
  .type  strcmp, @function
strcmp:
  or    a4, a0, a1
  li    t2, -1
  and   a4, a4, SZREG-1
  bnez  a4, .Lsimpleloop

  # Main loop for aligned strings
.Lloop:
  REG_L a2, 0(a0)
  REG_L a3, 0(a1)
  orc.b t0, a2
  bne   t0, t2, .Lfoundnull
  addi  a0, a0, SZREG
  addi  a1, a1, SZREG
  beq   a2, a3, .Lloop

  # Words don't match, and no null byte in first word.
  # Get bytes in big-endian order and compare.
#if __BYTE_ORDER__ == __ORDER_LITTLE_ENDIAN__
  rev8  a2, a2
  rev8  a3, a3
#endif
  # Synthesize (a2 >= a3) ? 1 : -1 in a branchless sequence.
  sltu a0, a2, a3
  neg  a0, a0
  ori  a0, a0, 1
  ret

.Lfoundnull:
  # Found a null byte.
  # If words don't match, fall back to simple loop.
  bne   a2, a3, .Lsimpleloop

  # Otherwise, strings are equal.
  li    a0, 0
  ret

  # Simple loop for misaligned strings
.Lsimpleloop:
  lbu   a2, 0(a0)
  lbu   a3, 0(a1)
  addi  a0, a0, 1
  addi  a1, a1, 1
  bne   a2, a3, 1f
  bnez  a2, .Lsimpleloop

1:
  sub   a0, a2, a3
  ret

.size   strcmp, .-strcmp

31. "P" Extension for Packed-SIMD Instructions, Version 0.2

Discussions at the 5th RISC-V workshop indicated a desire to drop this packed-SIMD proposal for floating-point registers in favor of standardizing on the V extension for large floating-point SIMD operations. However, there was interest in packed-SIMD fixed-point operations for use in the integer registers of small RISC-V implementations. A task group is working to define the new P extension.

32. "V" Standard Extension for Vector Operations, Version 1.0

CV32A65X: This extension is not supported.

33. Cryptography Extensions: Scalar & Entropy Source Instructions, Version 1.0.1

CV32A65X: These extensions are not supported.

34. Cryptography Extensions: Vector Instructions, Version 1.0

CV32A65X: These extensions are not supported.

35. Control-flow Integrity (CFI)

CV32A65X: The Zicfiss extension is not supported.

CV32A65X: The Zicfilp extension is not supported.

36. RV32/64G Instruction Set Listings

One goal of the RISC-V project is that it be used as a stable software development target. For this purpose, we define a combination of a base ISA (RV32I or RV64I) plus selected standard extensions (IMAFD, Zicsr, Zifencei) as a "general-purpose" ISA, and we use the abbreviation G for the IMAFDZicsr_Zifencei combination of instruction-set extensions.

CV32A65X: This chapter presents opcode maps and instruction-set listings for CV32A65X.

Table 26. RISC-V base opcode map, inst[1:0]=11

inst[4:2]

000

001

010

011

100

101

110

111 (>32b)

inst[6:5]

00

LOAD

LOAD-FP

custom-0

MISC-MEM

OP-IMM

AUIPC

OP-IMM-32

48b

01

STORE

STORE-FP

custom-1

AMO

OP

LUI

OP-32

64b

10

MADD

MSUB

NMSUB

NMADD

OP-FP

OP-V

custom-2/rv128

48b

11

BRANCH

JALR

reserved

JAL

SYSTEM

OP-VE

custom-3/rv128

≥80b

Table 26 shows a map of the major opcodes for RVG. Major opcodes with 3 or more lower bits set are reserved for instruction lengths greater than 32 bits. Opcodes marked as reserved should be avoided for custom instruction-set extensions as they might be used by future standard extensions. Major opcodes marked as custom-0 and custom-1 will be avoided by future standard extensions and are recommended for use by custom instruction-set extensions within the base 32-bit instruction format. The opcodes marked custom-2/rv128 and custom-3/rv128 are reserved for future use by RV128, but will otherwise be avoided for standard extensions and so can also be used for custom instruction-set extensions in RV32 and RV64.

We believe RV32G and RV64G provide simple but complete instruction sets for a broad range of general-purpose computing. The optional compressed instruction set described in Chapter 28 can be added (forming RV32GC and RV64GC) to improve performance, code size, and energy efficiency, though with some additional hardware complexity.

As we move beyond IMAFDC into further instruction-set extensions, the added instructions tend to be more domain-specific and only provide benefits to a restricted class of applications, e.g., for multimedia or security. Unlike most commercial ISAs, the RISC-V ISA design clearly separates the base ISA and broadly applicable standard extensions from these more specialized additions. Chapter 37 has a more extensive discussion of ways to add extensions to the RISC-V ISA.

31

27

26

25

24

20

19

15

14

12

11

7

6

0

funct7

rs2

rs1

funct3

rd

opcode

R-type

imm[11:0]

rs1

funct3

rd

opcode

I-type

imm[11:5]

rs2

rs1

funct3

imm[4:0]

opcode

S-type

imm[12|10:5]

rs2

rs1

funct3

imm[4:1|11]

opcode

B-type

imm[31:12]

rd

opcode

U-type

imm[20|10:1|11|19:12]

rd

opcode

J-type

RV32I Base Instruction Set

imm[31:12]

rd

0110111

LUI

imm[31:12]

rd

0010111

AUIPC

imm[20|10:1|11|19:12]

rd

1101111

JAL

imm[11:0]

rs1

000

rd

1100111

JALR

imm[12|10:5]

rs2

rs1

000

imm[4:1|11]

1100011

BEQ

imm[12|10:5]

rs2

rs1

001

imm[4:1|11]

1100011

BNE

imm[12|10:5]

rs2

rs1

100

imm[4:1|11]

1100011

BLT

imm[12|10:5]

rs2

rs1

101

imm[4:1|11]

1100011

BGE

imm[12|10:5]

rs2

rs1

110

imm[4:1|11]

1100011

BLTU

imm[12|10:5]

rs2

rs1

111

imm[4:1|11]

1100011

BGEU

imm[11:0]

rs1

000

rd

0000011

LB

imm[11:0]

rs1

001

rd

0000011

LH

imm[11:0]

rs1

010

rd

0000011

LW

imm[11:0]

rs1

100

rd

0000011

LBU

imm[11:0]

rs1

101

rd

0000011

LHU

imm[11:5]

rs2

rs1

000

imm[4:0]

0100011

SB

imm[11:5]

rs2

rs1

001

imm[4:0]

0100011

SH

imm[11:5]

rs2

rs1

010

imm[4:0]

0100011

SW

imm[11:0]

rs1

000

rd

0010011

ADDI

imm[11:0]

rs1

010

rd

0010011

SLTI

imm[11:0]

rs1

011

rd

0010011

SLTIU

imm[11:0]

rs1

100

rd

0010011

XORI

imm[11:0]

rs1

110

rd

0010011

ORI

imm[11:0]

rs1

111

rd

0010011

ANDI

0000000

shamt

rs1

001

rd

0010011

SLLI

0000000

shamt

rs1

101

rd

0010011

SRLI

0100000

shamt

rs1

101

rd

0010011

SRAI

0000000

rs2

rs1

000

rd

0110011

ADD

0100000

rs2

rs1

000

rd

0110011

SUB

0000000

rs2

rs1

001

rd

0110011

SLL

0000000

rs2

rs1

010

rd

0110011

SLT

0000000

rs2

rs1

011

rd

0110011

SLTU

0000000

rs2

rs1

100

rd

0110011

XOR

0000000

rs2

rs1

101

rd

0110011

SRL

0100000

rs2

rs1

101

rd

0110011

SRA

0000000

rs2

rs1

110

rd

0110011

OR

0000000

rs2

rs1

111

rd

0110011

AND

fm

pred

succ

rs1

000

rd

0001111

FENCE

1000

0011

0011

00000

000

00000

0001111

FENCE.TSO

0000

0001

0000

00000

000

00000

0001111

PAUSE

000000000000

00000

000

00000

1110011

ECALL

000000000001

00000

000

00000

1110011

EBREAK

31

27

26

25

24

20

19

15

14

12

11

7

6

0

funct7

rs2

rs1

funct3

rd

opcode

R-type

imm[11:0]

rs1

funct3

rd

opcode

I-type

imm[11:5]

rs2

rs1

funct3

imm[4:0]

opcode

S-type

RV32/RV64 Zicsr Standard Extension

csr

rs1

001

rd

1110011

CSRRW

csr

rs1

010

rd

1110011

CSRRS

csr

rs1

011

rd

1110011

CSRRC

csr

uimm

101

rd

1110011

CSRRWI

csr

uimm

110

rd

1110011

CSRRSI

csr

uimm

111

rd

1110011

CSRRCI

RV32M Standard Extension

0000001

rs2

rs1

000

rd

0110011

MUL

0000001

rs2

rs1

001

rd

0110011

MULH

0000001

rs2

rs1

010

rd

0110011

MULHSU

0000001

rs2

rs1

011

rd

0110011

MULHU

0000001

rs2

rs1

100

rd

0110011

DIV

0000001

rs2

rs1

101

rd

0110011

DIVU

0000001

rs2

rs1

110

rd

0110011

REM

0000001

rs2

rs1

111

rd

0110011

REMU

Table 27 lists the CSRs that have currently been allocated CSR addresses. The timers, counters, and floating-point CSRs are the only CSRs defined in this specification.

Table 27. RISC-V control and status register (CSR) address map.
Number Privilege Name Description

Floating-Point Control and Status Registers

0x001

Read write

fflags

Floating-Point Accrued Exceptions.

0x002

Read write

frm

Floating-Point Dynamic Rounding Mode.

0x003

Read write

fcsr

Floating-Point Control and Status Register (frm + fflags).

Counters and Timers

0xC00

Read-only

cycle

Cycle counter for RDCYCLE instruction.

0xC01

Read-only

time

Timer for RDTIME instruction.

0xC02

Read-only

instret

Instructions-retired counter for RDINSTRET instruction.

0xC80

Read-only

cycleh

Upper 32 bits of cycle, RV32I only.

0xC81

Read-only

timeh

Upper 32 bits of time, RV32I only.

0xC82

Read-only

instreth

Upper 32 bits of instret, RV32I only.

37. Extending RISC-V

In addition to supporting standard general-purpose software development, another goal of RISC-V is to provide a basis for more specialized instruction-set extensions or more customized accelerators. The instruction encoding spaces and optional variable-length instruction encoding are designed to make it easier to leverage software development effort for the standard ISA toolchain when building more customized processors. For example, the intent is to continue to provide full software support for implementations that only use the standard I base, perhaps together with many non-standard instruction-set extensions.

This chapter describes various ways in which the base RISC-V ISA can be extended, together with the scheme for managing instruction-set extensions developed by independent groups. This volume only deals with the unprivileged ISA, although the same approach and terminology is used for supervisor-level extensions described in the second volume.

37.1. Extension Terminology

This section defines some standard terminology for describing RISC-V extensions.

37.1.1. Standard versus Non-Standard Extension

Any RISC-V processor implementation must support a base integer ISA (RV32I, RV32E, RV64I, RV64E, or RV128I). In addition, an implementation may support one or more extensions. We divide extensions into two broad categories: standard versus non-standard.

  • A standard extension is one that is generally useful and that is designed to not conflict with any other standard extension. Currently, "MAFDQCBTPV", described in other chapters of this manual, are either complete or planned standard extensions.

  • A non-standard extension may be highly specialized and may conflict with other standard or non-standard extensions. We anticipate a wide variety of non-standard extensions will be developed over time, with some eventually being promoted to standard extensions.

37.1.2. Instruction Encoding Spaces and Prefixes

An instruction encoding space is some number of instruction bits within which a base ISA or ISA extension is encoded. RISC-V supports varying instruction lengths, but even within a single instruction length, there are various sizes of encoding space available. For example, the base ISAs are defined within a 30-bit encoding space (bits 31-2 of the 32-bit instruction), while the atomic extension "A" fits within a 25-bit encoding space (bits 31-7).

We use the term prefix to refer to the bits to the right of an instruction encoding space (since instruction fetch in RISC-V is little-endian, the bits to the right are stored at earlier memory addresses, hence form a prefix in instruction-fetch order). The prefix for the standard base ISA encoding is the two-bit "11" field held in bits 1-0 of the 32-bit word, while the prefix for the standard atomic extension "A" is the seven-bit "0101111" field held in bits 6-0 of the 32-bit word representing the AMO major opcode. A quirk of the encoding format is that the 3-bit funct3 field used to encode a minor opcode is not contiguous with the major opcode bits in the 32-bit instruction format, but is considered part of the prefix for 22-bit instruction spaces.

Although an instruction encoding space could be of any size, adopting a smaller set of common sizes simplifies packing independently developed extensions into a single global encoding. Table 28 gives the suggested sizes for RISC-V.

Table 28. Suggested standard RISC-V instruction encoding space sizes.
Size Usage # Available in standard instruction length

16-bit

32-bit

48-bit

64-bit

14-bit

Quadrant of compressed 16-bit encoding

3

22-bit

Minor opcode in base 32-bit encoding

stem c4baab0af7a156d93ff31ad113e2dde5

stem 3d91b2575aebc0d742415f31a7138964

stem a5ee56a7cd1f96b794dd9fd33c17b806

25-bit

Major opcode in base 32-bit encoding

32

stem cec3b648d9a79b7951e2288a4cb712cd

stem 4077d56d7f287ffbdb80f419a66226a9

30-bit

Quadrant of base 32-bit encoding

1

stem 6a484de5eab2c8d47af90cbe3ffedf6e

stem 840e9073a2fd95fa3f885a3e28972b40

32-bit

Minor opcode in 48-bit encoding

stem 560d002b3cbfbdc26102752336160ae5

stem 106e24e6dc15f353b281ecc18a55b796

37-bit

Major opcode in 48-bit encoding

32

stem 3d91b2575aebc0d742415f31a7138964

40-bit

Quadrant of 48-bit encoding

4

stem cec3b648d9a79b7951e2288a4cb712cd

45-bit

Sub-minor opcode in 64-bit encoding

stem 6a484de5eab2c8d47af90cbe3ffedf6e

48-bit

Minor opcode in 64-bit encoding

stem 96cee3dcfd74a2467659d99ac60808d7

52-bit

Major opcode in 64-bit encoding

32

37.1.3. Greenfield versus Brownfield Extensions

We use the term greenfield extension to describe an extension that begins populating a new instruction encoding space, and hence can only cause encoding conflicts at the prefix level. We use the term brownfield extension to describe an extension that fits around existing encodings in a previously defined instruction space. A brownfield extension is necessarily tied to a particular greenfield parent encoding, and there may be multiple brownfield extensions to the same greenfield parent encoding. For example, the base ISAs are greenfield encodings of a 30-bit instruction space, while the FDQ floating-point extensions are all brownfield extensions adding to the parent base ISA 30-bit encoding space.

Note that we consider the standard A extension to have a greenfield encoding as it defines a new previously empty 25-bit encoding space in the leftmost bits of the full 32-bit base instruction encoding, even though its standard prefix locates it within the 30-bit encoding space of its parent base ISA. Changing only its single 7-bit prefix could move the A extension to a different 30-bit encoding space while only worrying about conflicts at the prefix level, not within the encoding space itself.

Table 29. Two-dimensional characterization of standard instruction-set extensions.
Adds state No new state

Greenfield

RV32I(30), RV64I(30)

A(25)

Brownfield

F(I), D(F), Q(D)

M(I)

Table 29 shows the bases and standard extensions placed in a simple two-dimensional taxonomy. One axis is whether the extension is greenfield or brownfield, while the other axis is whether the extension adds architectural state. For greenfield extensions, the size of the instruction encoding space is given in parentheses. For brownfield extensions, the name of the extension (greenfield or brownfield) it builds upon is given in parentheses. Additional user-level architectural state usually implies changes to the supervisor-level system or possibly to the standard calling convention.

Note that RV64I is not considered an extension of RV32I, but a different complete base encoding.

37.1.4. Standard-Compatible Global Encodings

A complete or global encoding of an ISA for an actual RISC-V implementation must allocate a unique non-conflicting prefix for every included instruction encoding space. The bases and every standard extension have each had a standard prefix allocated to ensure they can all coexist in a global encoding.

A standard-compatible global encoding is one where the base and every included standard extension have their standard prefixes. A standard-compatible global encoding can include non-standard extensions that do not conflict with the included standard extensions. A standard-compatible global encoding can also use standard prefixes for non-standard extensions if the associated standard extensions are not included in the global encoding. In other words, a standard extension must use its standard prefix if included in a standard-compatible global encoding, but otherwise its prefix is free to be reallocated. These constraints allow a common toolchain to target the standard subset of any RISC-V standard-compatible global encoding.

37.1.5. Guaranteed Non-Standard Encoding Space

To support development of proprietary custom extensions, portions of the encoding space are guaranteed to never be used by standard extensions.

37.2. RISC-V Extension Design Philosophy

We intend to support a large number of independently developed extensions by encouraging extension developers to operate within instruction encoding spaces, and by providing tools to pack these into a standard-compatible global encoding by allocating unique prefixes. Some extensions are more naturally implemented as brownfield augmentations of existing extensions, and will share whatever prefix is allocated to their parent greenfield extension. The standard extension prefixes avoid spurious incompatibilities in the encoding of core functionality, while allowing custom packing of more esoteric extensions.

This capability of repacking RISC-V extensions into different standard-compatible global encodings can be used in a number of ways.

One use-case is developing highly specialized custom accelerators, designed to run kernels from important application domains. These might want to drop all but the base integer ISA and add in only the extensions that are required for the task in hand. The base ISAs have been designed to place minimal requirements on a hardware implementation, and has been encoded to use only a small fraction of a 32-bit instruction encoding space.

Another use-case is to build a research prototype for a new type of instruction-set extension. The researchers might not want to expend the effort to implement a variable-length instruction-fetch unit, and so would like to prototype their extension using a simple 32-bit fixed-width instruction encoding. However, this new extension might be too large to coexist with standard extensions in the 32-bit space. If the research experiments do not need all of the standard extensions, a standard-compatible global encoding might drop the unused standard extensions and reuse their prefixes to place the proposed extension in a non-standard location to simplify engineering of the research prototype. Standard tools will still be able to target the base and any standard extensions that are present to reduce development time. Once the instruction-set extension has been evaluated and refined, it could then be made available for packing into a larger variable-length encoding space to avoid conflicts with all standard extensions.

The following sections describe increasingly sophisticated strategies for developing implementations with new instruction-set extensions. These are mostly intended for use in highly customized, educational, or experimental architectures rather than for the main line of RISC-V ISA development.

37.3. Extensions within fixed-width 32-bit instruction format

In this section, we discuss adding extensions to implementations that only support the base fixed-width 32-bit instruction format.

We anticipate the simplest fixed-width 32-bit encoding will be popular for many restricted accelerators and research prototypes.

37.3.1. Available 30-bit instruction encoding spaces

In the standard encoding, three of the available 30-bit instruction encoding spaces (those with 2-bit prefixes 00, 01, and 10) are used to enable the optional compressed instruction extension. However, if the compressed instruction-set extension is not required, then these three further 30-bit encoding spaces become available. This quadruples the available encoding space within the 32-bit format.

37.3.2. Available 25-bit instruction encoding spaces

A 25-bit instruction encoding space corresponds to a major opcode in the base and standard extension encodings.

There are four major opcodes expressly designated for custom extensions Table 26, each of which represents a 25-bit encoding space. Two of these are reserved for eventual use in the RV128 base encoding (will be OP-IMM-64 and OP-64), but can be used for non-standard extensions for RV32 and RV64.

The two major opcodes reserved for RV64 (OP-IMM-32 and OP-32) can also be used for non-standard extensions to RV32 only.

If an implementation does not require floating-point, then the seven major opcodes reserved for standard floating-point extensions (LOAD-FP, STORE-FP, MADD, MSUB, NMSUB, NMADD, OP-FP) can be reused for non-standard extensions. Similarly, the AMO major opcode can be reused if the standard atomic extensions are not required.

If an implementation does not require instructions longer than 32-bits, then an additional four major opcodes are available (those marked in gray in Table 26).

The base RV32I encoding uses only 11 major opcodes plus 3 reserved opcodes, leaving up to 18 available for extensions. The base RV64I encoding uses only 13 major opcodes plus 3 reserved opcodes, leaving up to 16 available for extensions.

37.3.3. Available 22-bit instruction encoding spaces

A 22-bit encoding space corresponds to a funct3 minor opcode space in the base and standard extension encodings. Several major opcodes have a funct3 field minor opcode that is not completely occupied, leaving available several 22-bit encoding spaces.

Usually a major opcode selects the format used to encode operands in the remaining bits of the instruction, and ideally, an extension should follow the operand format of the major opcode to simplify hardware decoding.

37.3.4. Other spaces

Smaller spaces are available under certain major opcodes, and not all minor opcodes are entirely filled.

37.4. Adding aligned 64-bit instruction extensions

The simplest approach to provide space for extensions that are too large for the base 32-bit fixed-width instruction format is to add naturally aligned 64-bit instructions. The implementation must still support the 32-bit base instruction format, but can require that 64-bit instructions are aligned on 64-bit boundaries to simplify instruction fetch, with a 32-bit NOP instruction used as alignment padding where necessary.

To simplify use of standard tools, the 64-bit instructions should be encoded as described in Table 1. However, an implementation might choose a non-standard instruction-length encoding for 64-bit instructions, while retaining the standard encoding for 32-bit instructions. For example, if compressed instructions are not required, then a 64-bit instruction could be encoded using one or more zero bits in the first two bits of an instruction.

We anticipate processor generators that produce instruction-fetch units capable of automatically handling any combination of supported variable-length instruction encodings.

37.5. Supporting VLIW encodings

Although RISC-V was not designed as a base for a pure VLIW machine, VLIW encodings can be added as extensions using several alternative approaches. In all cases, the base 32-bit encoding has to be supported to allow use of any standard software tools.

37.5.1. Fixed-size instruction group

The simplest approach is to define a single large naturally aligned instruction format (e.g., 128 bits) within which VLIW operations are encoded. In a conventional VLIW, this approach would tend to waste instruction memory to hold NOPs, but a RISC-V-compatible implementation would have to also support the base 32-bit instructions, confining the VLIW code size expansion to VLIW-accelerated functions.

37.5.2. Encoded-Length Groups

Another approach is to use the standard length encoding from Table 1 to encode parallel instruction groups, allowing NOPs to be compressed out of the VLIW instruction. For example, a 64-bit instruction could hold two 28-bit operations, while a 96-bit instruction could hold three 28-bit operations, and so on. Alternatively, a 48-bit instruction could hold one 42-bit operation, while a 96-bit instruction could hold two 42-bit operations, and so on.

This approach has the advantage of retaining the base ISA encoding for instructions holding a single operation, but has the disadvantage of requiring a new 28-bit or 42-bit encoding for operations within the VLIW instructions, and misaligned instruction fetch for larger groups. One simplification is to not allow VLIW instructions to straddle certain microarchitecturally significant boundaries (e.g., cache lines or virtual memory pages).

37.5.3. Fixed-Size Instruction Bundles

Another approach, similar to Itanium, is to use a larger naturally aligned fixed instruction bundle size (e.g., 128 bits) across which parallel operation groups are encoded. This simplifies instruction fetch, but shifts the complexity to the group execution engine. To remain RISC-V compatible, the base 32-bit instruction would still have to be supported.

37.5.4. End-of-Group bits in Prefix

None of the above approaches retains the RISC-V encoding for the individual operations within a VLIW instruction. Yet another approach is to repurpose the two prefix bits in the fixed-width 32-bit encoding. One prefix bit can be used to signal "end-of-group" if set, while the second bit could indicate execution under a predicate if clear. Standard RISC-V 32-bit instructions generated by tools unaware of the VLIW extension would have both prefix bits set (11) and thus have the correct semantics, with each instruction at the end of a group and not predicated.

The main disadvantage of this approach is that the base ISAs lack the complex predication support usually required in an aggressive VLIW system, and it is difficult to add space to specify more predicate registers in the standard 30-bit encoding space.

38. ISA Extension Naming Conventions

This chapter describes the RISC-V ISA extension naming scheme that is used to concisely describe the set of instructions present in a hardware implementation, or the set of instructions used by an application binary interface (ABI).

The RISC-V ISA is designed to support a wide variety of implementations with various experimental instruction-set extensions. We have found that an organized naming scheme simplifies software tools and documentation.

38.1. Case Sensitivity

The ISA naming strings are case insensitive.

38.2. Base Integer ISA

RISC-V ISA strings begin with either RV32I, RV32E, RV64I, RV64E, or RV128I indicating the supported address space size in bits for the base integer ISA.

38.3. Instruction-Set Extension Names

Standard ISA extensions are given a name consisting of a single letter. For example, the first four standard extensions to the integer bases are: "M" for integer multiplication and division, "A" for atomic memory instructions, "F" for single-precision floating-point instructions, and "D" for double-precision floating-point instructions. Any RISC-V instruction-set variant can be succinctly described by concatenating the base integer prefix with the names of the included extensions, e.g., "RV64IMAFD".

We have also defined an abbreviation "G" to represent the "IMAFDZicsr_Zifencei" base and extensions, as this is intended to represent our standard general-purpose ISA.

Standard extensions to the RISC-V ISA are given other reserved letters, e.g., "Q" for quad-precision floating-point, or "C" for the 16-bit compressed instruction format.

Some ISA extensions depend on the presence of other extensions, e.g., "D" depends on "F" and "F" depends on "Zicsr". These dependencies may be implicit in the ISA name: for example, RV32IF is equivalent to RV32IFZicsr, and RV32ID is equivalent to RV32IFD and RV32IFDZicsr.

38.4. Underscores

Underscores "_" may be used to separate ISA extensions to improve readability and to provide disambiguation, e.g., "RV32I2_M2_A2".

38.5. Additional Standard Unprivileged Extension Names

Standard unprivileged extensions can also be named by using a single "Z" followed by an alphanumeric name. The name must end with an alphabetical character. The second letter from the end cannot be numeric if the last letter is "p". For example, "Zifencei" names the instruction-fetch fence extension described in Chapter 6.

The first letter following the "Z" conventionally indicates the most closely related alphabetical extension category, IMAFDQLCBKJTPVH. For the "Zfa" extension for additional floating-point instructions, for example, the letter "f" indicates the extension is related to the "F" standard extension. If multiple "Z" extensions are named, they should be ordered first by category, then alphabetically within a category—for example, "Zicsr_Zifencei_Ztso".

All multi-letter extensions, including those with the "Z" prefix, must be separated from other multi-letter extensions by an underscore, e.g., "RV32IMACZicsr_Zifencei".

38.6. Supervisor-level Instruction-Set Extension Names

Standard extensions that extend the supervisor-level virtual-memory architecture are prefixed with the letters "Sv", followed by an alphanumeric name. Other standard extensions that extend the supervisor-level architecture are prefixed with thel letters "Ss", followed by an alphanumeric name. The name must end with an alphabetical character. The second letter from the end cannot be numeric if the last letter is "p". These extensions are further defined in Volume II.

The extensions "sv32", "sv39", "sv48", and "sv59" were defined before the rule against extension names ending in numbers was established.

Standard supervisor-level extensions should be listed after standard unprivileged extensions, and like other multi-letter extensions, must be separated from other multi-letter extensions by an underscore. If multiple supervisor-level extensions are listed, they should be ordered alphabetically.

38.7. Hypervisor-level Instruction-Set Extension Names

Standard extensions that extend the hypervisor-level architecture are prefixed with the letters "Sh". If multiple hypervisor-level extensions are listed, they should be ordered alphabetically.

Many augmentations to the hypervisor-level architecture are more naturally defined as supervisor-level extensions, following the scheme described in the previous section. The "Sh" prefix is used by the few hypervisor-level extensions that have no supervisor-visible effects.

38.8. Machine-level Instruction-Set Extension Names

Standard machine-level instruction-set extensions are prefixed with the letters "Sm".

Standard machine-level extensions should be listed after standard lesser-privileged extensions, and like other multi-letter extensions, must be separated from other multi-letter extensions by an underscore. If multiple machine-level extensions are listed, they should be ordered alphabetically.

38.9. Non-Standard Extension Names

Non-standard extensions are named by using a single "X" followed by the alphanumeric name. The name must end with an alphabetic character. The second letter from the end cannot be numeric if the last letter is "p". For example, "Xhwacha" names the Hwacha vector-fetch ISA extension.

Non-standard extensions must be listed after all standard extensions, and, like other multi-letter extensions, must be separated from other multi-letter extensions by an underscore. For example, an ISA with non-standard extensions Argle and Bargle may be named "RV64IZifencei_Xargle_Xbargle".

If multiple non-standard extensions are listed, they should be ordered alphabetically. Like other multi-letter extensions, they should be separated from other multi-leter extensions by an underscore.

38.10. Version Numbers

Recognizing that instruction sets may expand or alter over time, we encode extension version numbers following the extension name. Version numbers are divided into major and minor version numbers, separated by a "p". If the minor version is "0", then "p0" can be omitted from the version string. To avoid ambiguity, no extension name may end with a number or a "p" preceded by a number.

Because the "P" extension for Packed SIMD can be confused for the decimal point in a version number, it must be preceded by an underscore if it follows another extension with a version number. For example, "rv32i2p2" means version 2.2 of RV32I, whereas "rv32i2_p2" means version 2.0 of RV32I with version 2.0 of the P extension.

Changes in major version numbers imply a loss of backwards compatibility, whereas changes in only the minor version number must be backwards-compatible. For example, the original 64-bit standard ISA defined in release 1.0 of this manual can be written in full as "RV64I1p0M1p0A1p0F1p0D1p0", more concisely as "RV64I1M1A1F1D1".

We introduced the version numbering scheme with the second release. Hence, we define the default version of a standard extension to be the version present at that time, e.g., "RV32I" is equivalent to "RV32I2".

38.11. Subset Naming Convention

Table 30 summarizes the standardized extension names. The table also defines the canonical order in which extension names must appear in the name string, with top-to-bottom in table indicating first-to-last in the name string, e.g., RV32IMACV is legal, whereas RV32IMAVC is not.

Table 30. Standard ISA extension names.
Subset Name Implies

Base ISA

Integer

I

Reduced Integer

E

Standard Unprivileged Extensions

Integer Multiplication and Division

M

Zmmul

Atomics

A

Single-Precision Floating-Point

F

Zicsr

Double-Precision Floating-Point

D

F

General

G

IMAFDZicsr_Zifencei

Quad-Precision Floating-Point

Q

D

16-bit Compressed Instructions

C

B Extension

B

Packed-SIMD Extensions

P

Vector Extension

V

D

Hypervisor Extension

H

Additional Standard Unprivileged Extensions

Additional Standard unprivileged extensions "abc"

Zabc

Standard Supervisor-Level Extensions

Supervisor-level extension "def"

Ssdef

Standard Hypervisor-Level Extensions

Hypervisor-level extension "ghi"

Shghi

Standard Machine-Level Extensions

Machine-level extension "jkl"

Smjkl

Non-Standard Extensions

Non-standard extension "mno"

Xmno

39. History and Acknowledgments

39.1. "Why Develop a new ISA?" Rationale from Berkeley Group

We developed RISC-V to support our own needs in research and education, where our group is particularly interested in actual hardware implementations of research ideas (we have completed eleven different silicon fabrications of RISC-V since the first edition of this specification), and in providing real implementations for students to explore in classes (RISC-V processor RTL designs have been used in multiple undergraduate and graduate classes at Berkeley). In our current research, we are especially interested in the move towards specialized and heterogeneous accelerators, driven by the power constraints imposed by the end of conventional transistor scaling. We wanted a highly flexible and extensible base ISA around which to build our research effort.

A question we have been repeatedly asked is "Why develop a new ISA?" The biggest obvious benefit of using an existing commercial ISA is the large and widely supported software ecosystem, both development tools and ported applications, which can be leveraged in research and teaching. Other benefits include the existence of large amounts of documentation and tutorial examples. However, our experience of using commercial instruction sets for research and teaching is that these benefits are smaller in practice, and do not outweigh the disadvantages:

  • Commercial ISAs are proprietary. Except for SPARC V8, which is an open IEEE standard (IEEE Standard for a 32-Bit Microprocessor, 1994) , most owners of commercial ISAs carefully guard their intellectual property and do not welcome freely available competitive implementations. This is much less of an issue for academic research and teaching using only software simulators, but has been a major concern for groups wishing to share actual RTL implementations. It is also a major concern for entities who do not want to trust the few sources of commercial ISA implementations, but who are prohibited from creating their own clean room implementations. We cannot guarantee that all RISC-V implementations will be free of third-party patent infringements, but we can guarantee we will not attempt to sue a RISC-V implementor.

  • Commercial ISAs are only popular in certain market domains. The most obvious examples at time of writing are that the ARM architecture is not well supported in the server space, and the Intel x86 architecture (or for that matter, almost every other architecture) is not well supported in the mobile space, though both Intel and ARM are attempting to enter each other’s market segments. Another example is ARC and Tensilica, which provide extensible cores but are focused on the embedded space. This market segmentation dilutes the benefit of supporting a particular commercial ISA as in practice the software ecosystem only exists for certain domains, and has to be built for others.

  • Commercial ISAs come and go. Previous research infrastructures have been built around commercial ISAs that are no longer popular (SPARC, MIPS) or even no longer in production (Alpha). These lose the benefit of an active software ecosystem, and the lingering intellectual property issues around the ISA and supporting tools interfere with the ability of interested third parties to continue supporting the ISA. An open ISA might also lose popularity, but any interested party can continue using and developing the ecosystem.

  • Popular commercial ISAs are complex. The dominant commercial ISAs (x86 and ARM) are both very complex to implement in hardware to the level of supporting common software stacks and operating systems. Worse, nearly all the complexity is due to bad, or at least outdated, ISA design decisions rather than features that truly improve efficiency.

  • Commercial ISAs alone are not enough to bring up applications. Even if we expend the effort to implement a commercial ISA, this is not enough to run existing applications for that ISA. Most applications need a complete ABI (application binary interface) to run, not just the user-level ISA. Most ABIs rely on libraries, which in turn rely on operating system support. To run an existing operating system requires implementing the supervisor-level ISA and device interfaces expected by the OS. These are usually much less well-specified and considerably more complex to implement than the user-level ISA.

  • Popular commercial ISAs were not designed for extensibility. The dominant commercial ISAs were not particularly designed for extensibility, and as a consequence have added considerable instruction encoding complexity as their instruction sets have grown. Companies such as Tensilica (acquired by Cadence) and ARC (acquired by Synopsys) have built ISAs and toolchains around extensibility, but have focused on embedded applications rather than general-purpose computing systems.

  • A modified commercial ISA is a new ISA. One of our main goals is to support architecture research, including major ISA extensions. Even small extensions diminish the benefit of using a standard ISA, as compilers have to be modified and applications rebuilt from source code to use the extension. Larger extensions that introduce new architectural state also require modifications to the operating system. Ultimately, the modified commercial ISA becomes a new ISA, but carries along all the legacy baggage of the base ISA.

Our position is that the ISA is perhaps the most important interface in a computing system, and there is no reason that such an important interface should be proprietary. The dominant commercial ISAs are based on instruction-set concepts that were already well known over 30 years ago. Software developers should be able to target an open standard hardware target, and commercial processor designers should compete on implementation quality.

We are far from the first to contemplate an open ISA design suitable for hardware implementation. We also considered other existing open ISA designs, of which the closest to our goals was the OpenRISC architecture (OpenCores, 2012). We decided against adopting the OpenRISC ISA for several technical reasons:

  • OpenRISC has condition codes and branch delay slots, which complicate higher performance implementations.

  • OpenRISC uses a fixed 32-bit encoding and 16-bit immediates, which precludes a denser instruction encoding and limits space for later expansion of the ISA.

  • OpenRISC does not support the 2008 revision to the IEEE 754 floating-point standard.

  • The OpenRISC 64-bit design had not been completed when we began.

By starting from a clean slate, we could design an ISA that met all of our goals, though of course, this took far more effort than we had planned at the outset. We have now invested considerable effort in building up the RISC-V ISA infrastructure, including documentation, compiler tool chains, operating system ports, reference ISA simulators, FPGA implementations, efficient ASIC implementations, architecture test suites, and teaching materials. Since the last edition of this manual, there has been considerable uptake of the RISC-V ISA in both academia and industry, and we have created the non-profit RISC-V Foundation to protect and promote the standard. The RISC-V Foundation website at riscv.org contains the latest information on the Foundation membership and various open-source projects using RISC-V.

39.2. History from Revision 1.0 of ISA manual

The RISC-V ISA and instruction-set manual builds upon several earlier projects. Several aspects of the supervisor-level machine and the overall format of the manual date back to the T0 (Torrent-0) vector microprocessor project at UC Berkeley and ICSI, begun in 1992. T0 was a vector processor based on the MIPS-II ISA, with Krste Asanović as main architect and RTL designer, and Brian Kingsbury and Bertrand Irrisou as principal VLSI implementors. David Johnson at ICSI was a major contributor to the T0 ISA design, particularly supervisor mode, and to the manual text. John Hauser also provided considerable feedback on the T0 ISA design.

The Scale (Software-Controlled Architecture for Low Energy) project at MIT, begun in 2000, built upon the T0 project infrastructure, refined the supervisor-level interface, and moved away from the MIPS scalar ISA by dropping the branch delay slot. Ronny Krashinsky and Christopher Batten were the principal architects of the Scale Vector-Thread processor at MIT, while Mark Hampton ported the GCC-based compiler infrastructure and tools for Scale.

A lightly edited version of the T0 MIPS scalar processor specification (MIPS-6371) was used in teaching a new version of the MIT 6.371 Introduction to VLSI Systems class in the Fall 2002 semester, with Chris Terman and Krste Asanović as lecturers. Chris Terman contributed most of the lab material for the class (there was no TA!). The 6.371 class evolved into the trial 6.884 Complex Digital Design class at MIT, taught by Arvind and Krste Asanović in Spring 2005, which became a regular Spring class 6.375. A reduced version of the Scale MIPS-based scalar ISA, named SMIPS, was used in 6.884/6.375. Christopher Batten was the TA for the early offerings of these classes and developed a considerable amount of documentation and lab material based around the SMIPS ISA. This same SMIPS lab material was adapted and enhanced by TA Yunsup Lee for the UC Berkeley Fall 2009 CS250 VLSI Systems Design class taught by John Wawrzynek, Krste Asanović, and John Lazzaro.

The Maven (Malleable Array of Vector-thread ENgines) project was a second-generation vector-thread architecture. Its design was led by Christopher Batten when he was an Exchange Scholar at UC Berkeley starting in summer 2007. Hidetaka Aoki, a visiting industrial fellow from Hitachi, gave considerable feedback on the early Maven ISA and microarchitecture design. The Maven infrastructure was based on the Scale infrastructure but the Maven ISA moved further away from the MIPS ISA variant defined in Scale, with a unified floating-point and integer register file. Maven was designed to support experimentation with alternative data-parallel accelerators. Yunsup Lee was the main implementor of the various Maven vector units, while Rimas Avižienis was the main implementor of the various Maven scalar units. Yunsup Lee and Christopher Batten ported GCC to work with the new Maven ISA. Christopher Celio provided the initial definition of a traditional vector instruction set ("Flood") variant of Maven.

Based on experience with all these previous projects, the RISC-V ISA definition was begun in Summer 2010, with Andrew Waterman, Yunsup Lee, Krste Asanović, and David Patterson as principal designers. An initial version of the RISC-V 32-bit instruction subset was used in the UC Berkeley Fall 2010 CS250 VLSI Systems Design class, with Yunsup Lee as TA. RISC-V is a clean break from the earlier MIPS-inspired designs. John Hauser contributed to the floating-point ISA definition, including the sign-injection instructions and a register encoding scheme that permits internal recoding of floating-point values.

39.3. History from Revision 2.0 of ISA manual

Multiple implementations of RISC-V processors have been completed, including several silicon fabrications, as shown in Fabricated RISC-V testchips table.

Name Tapeout Date Process ISA

Raven-1

May 29, 2011

ST 28nm FDSOI

RV64G1_Xhwacha1

EOS14

April 1, 2012

IBM 45nm SOI

RV64G1p1_Xhwacha2

EOS16

August 17, 2012

IBM 45nm SOI

RV64G1p1_Xhwacha2

Raven-2

August 22, 2012

ST 28nm FDSOI

RV64G1p1_Xhwacha2

EOS18

February 6, 2013

IBM 45nm SOI

RV64G1p1_Xhwacha2

EOS20

July 3, 2013

IBM 45nm SOI

RV64G1p99_Xhwacha2

Raven-3

September 26, 2013

ST 28nm SOI

RV64G1p99_Xhwacha2

EOS22

March 7, 2014

IBM 45nm SOI

RV64G1p9999_Xhwacha3

The first RISC-V processors to be fabricated were written in Verilog and manufactured in a pre-production FDSOI technology from ST as the Raven-1 testchip in 2011. Two cores were developed by Yunsup Lee and Andrew Waterman, advised by Krste Asanović, and fabricated together: 1) an RV64 scalar core with error-detecting flip-flops, and 2) an RV64 core with an attached 64-bit floating-point vector unit. The first microarchitecture was informally known as "TrainWreck", due to the short time available to complete the design with immature design libraries.

Subsequently, a clean microarchitecture for an in-order decoupled RV64 core was developed by Andrew Waterman, Rimas Avižienis, and Yunsup Lee, advised by Krste Asanović, and, continuing the railway theme, was codenamed "Rocket" after George Stephenson’s successful steam locomotive design. Rocket was written in Chisel, a new hardware design language developed at UC Berkeley. The IEEE floating-point units used in Rocket were developed by John Hauser, Andrew Waterman, and Brian Richards. Rocket has since been refined and developed further, and has been fabricated two more times in FDSOI (Raven-2, Raven-3), and five times in IBM SOI technology (EOS14, EOS16, EOS18, EOS20, EOS22) for a photonics project. Work is ongoing to make the Rocket design available as a parameterized RISC-V processor generator.

EOS14-EOS22 chips include early versions of Hwacha, a 64-bit IEEE floating-point vector unit, developed by Yunsup Lee, Andrew Waterman, Huy Vo, Albert Ou, Quan Nguyen, and Stephen Twigg, advised by Krste Asanović. EOS16-EOS22 chips include dual cores with a cache-coherence protocol developed by Henry Cook and Andrew Waterman, advised by Krste Asanović. EOS14 silicon has successfully run at 1.25 GHz. EOS16 silicon suffered from a bug in the IBM pad libraries. EOS18 and EOS20 have successfully run at 1.35 GHz.

Contributors to the Raven testchips include Yunsup Lee, Andrew Waterman, Rimas Avižienis, Brian Zimmer, Jaehwa Kwak, Ruzica Jevtić, Milovan Blagojević, Alberto Puggelli, Steven Bailey, Ben Keller, Pi-Feng Chiu, Brian Richards, Borivoje Nikolić, and Krste Asanović.

Contributors to the EOS testchips include Yunsup Lee, Rimas Avižienis, Andrew Waterman, Henry Cook, Huy Vo, Daiwei Li, Chen Sun, Albert Ou, Quan Nguyen, Stephen Twigg, Vladimir Stojanović, and Krste Asanović.

Andrew Waterman and Yunsup Lee developed the C++ ISA simulator "Spike", used as a golden model in development and named after the golden spike used to celebrate completion of the US transcontinental railway. Spike has been made available as a BSD open-source project.

Andrew Waterman completed a Master’s thesis with a preliminary design of the RISC-V compressed instruction set (Waterman, 2011).

Various FPGA implementations of the RISC-V have been completed, primarily as part of integrated demos for the Par Lab project research retreats. The largest FPGA design has 3 cache-coherent RV64IMA processors running a research operating system. Contributors to the FPGA implementations include Andrew Waterman, Yunsup Lee, Rimas Avižienis, and Krste Asanović.

RISC-V processors have been used in several classes at UC Berkeley. Rocket was used in the Fall 2011 offering of CS250 as a basis for class projects, with Brian Zimmer as TA. For the undergraduate CS152 class in Spring 2012, Christopher Celio used Chisel to write a suite of educational RV32 processors, named "Sodor" after the island on which "Thomas the Tank Engine" and friends live. The suite includes a microcoded core, an unpipelined core, and 2, 3, and 5-stage pipelined cores, and is publicly available under a BSD license. The suite was subsequently updated and used again in CS152 in Spring 2013, with Yunsup Lee as TA, and in Spring 2014, with Eric Love as TA. Christopher Celio also developed an out-of-order RV64 design known as BOOM (Berkeley Out-of-Order Machine), with accompanying pipeline visualizations, that was used in the CS152 classes. The CS152 classes also used cache-coherent versions of the Rocket core developed by Andrew Waterman and Henry Cook.

Over the summer of 2013, the RoCC (Rocket Custom Coprocessor) interface was defined to simplify adding custom accelerators to the Rocket core. Rocket and the RoCC interface were used extensively in the Fall 2013 CS250 VLSI class taught by Jonathan Bachrach, with several student accelerator projects built to the RoCC interface. The Hwacha vector unit has been rewritten as a RoCC coprocessor.

Two Berkeley undergraduates, Quan Nguyen and Albert Ou, have successfully ported Linux to run on RISC-V in Spring 2013.

Colin Schmidt successfully completed an LLVM backend for RISC-V 2.0 in January 2014.

Darius Rad at Bluespec contributed soft-float ABI support to the GCC port in March 2014.

John Hauser contributed the definition of the floating-point classification instructions.

We are aware of several other RISC-V core implementations, including one in Verilog by Tommy Thorn, and one in Bluespec by Rishiyur Nikhil.

39.4. Acknowledgments

Thanks to Christopher F. Batten, Preston Briggs, Christopher Celio, David Chisnall, Stefan Freudenberger, John Hauser, Ben Keller, Rishiyur Nikhil, Michael Taylor, Tommy Thorn, and Robert Watson for comments on the draft ISA version 2.0 specification.

39.5. History from Revision 2.1

Uptake of the RISC-V ISA has been very rapid since the introduction of the frozen version 2.0 in May 2014, with too much activity to record in a short history section such as this. Perhaps the most important single event was the formation of the non-profit RISC-V Foundation in August 2015. The Foundation will now take over stewardship of the official RISC-V ISA standard, and the official website riscv.org is the best place to obtain news and updates on the RISC-V standard.

39.6. Acknowledgments

Thanks to Scott Beamer, Allen J. Baum, Christopher Celio, David Chisnall, Paul Clayton, Palmer Dabbelt, Jan Gray, Michael Hamburg, and John Hauser for comments on the version 2.0 specification.

39.7. History from Revision 2.2

39.8. Acknowledgments

Thanks to Jacob Bachmeyer, Alex Bradbury, David Horner, Stefan O’Rear, and Joseph Myers for comments on the version 2.1 specification.

39.9. History for Revision 2.3

Uptake of RISC-V continues at a breakneck pace.

John Hauser and Andrew Waterman contributed a hypervisor ISA extension based upon a proposal from Paolo Bonzini.

Daniel Lustig, Arvind, Krste Asanović, Shaked Flur, Paul Loewenstein, Yatin Manerkar, Luc Maranget, Margaret Martonosi, Vijayanand Nagarajan, Rishiyur Nikhil, Jonas Oberhauser, Christopher Pulte, Jose Renau, Peter Sewell, Susmit Sarkar, Caroline Trippel, Muralidaran Vijayaraghavan, Andrew Waterman, Derek Williams, Andrew Wright, and Sizhuo Zhang contributed the memory consistency model.

39.10. Funding

Development of the RISC-V architecture and implementations has been partially funded by the following sponsors.

  • Par Lab: Research supported by Microsoft (Award # 024263) and Intel (Award # 024894) funding and by matching funding by U.C. Discovery (Award # DIG07-10227). Additional support came from Par Lab affiliates Nokia, NVIDIA, Oracle, and Samsung.

  • Project Isis: DoE Award DE-SC0003624.

  • ASPIRE Lab: DARPA PERFECT program, Award HR0011-12-2-0016. DARPA POEM program Award HR0011-11-C-0100. The Center for Future Architectures Research (C-FAR), a STARnet center funded by the Semiconductor Research Corporation. Additional support from ASPIRE industrial sponsor, Intel, and ASPIRE affiliates, Google, Hewlett Packard Enterprise, Huawei, Nokia, NVIDIA, Oracle, and Samsung.

The content of this paper does not necessarily reflect the position or the policy of the US government and no official endorsement should be inferred.

Appendix A: RVWMO Explanatory Material, Version 0.1

This section provides more explanation for RVWMO Chapter 18, using more informal language and concrete examples. These are intended to clarify the meaning and intent of the axioms and preserved program order rules. This appendix should be treated as commentary; all normative material is provided in Chapter 18 and in the rest of the main body of the ISA specification. All currently known discrepancies are listed in Section A.7. Any other discrepancies are unintentional.

A.1. Why RVWMO?

Memory consistency models fall along a loose spectrum from weak to strong. Weak memory models allow more hardware implementation flexibility and deliver arguably better performance, performance per watt, power, scalability, and hardware verification overheads than strong models, at the expense of a more complex programming model. Strong models provide simpler programming models, but at the cost of imposing more restrictions on the kinds of (non-speculative) hardware optimizations that can be performed in the pipeline and in the memory system, and in turn imposing some cost in terms of power, area overhead, and verification burden.

RISC-V has chosen the RVWMO memory model, a variant of release consistency. This places it in between the two extremes of the memory model spectrum. The RVWMO memory model enables architects to build simple implementations, aggressive implementations, implementations embedded deeply inside a much larger system and subject to complex memory system interactions, or any number of other possibilities, all while simultaneously being strong enough to support programming language memory models at high performance.

To facilitate the porting of code from other architectures, some hardware implementations may choose to implement the Ztso extension, which provides stricter RVTSO ordering semantics by default. Code written for RVWMO is automatically and inherently compatible with RVTSO, but code written assuming RVTSO is not guaranteed to run correctly on RVWMO implementations. In fact, most RVWMO implementations will (and should) simply refuse to run RVTSO-only binaries. Each implementation must therefore choose whether to prioritize compatibility with RVTSO code (e.g., to facilitate porting from x86) or whether to instead prioritize compatibility with other RISC-V cores implementing RVWMO.

Some fences and/or memory ordering annotations in code written for RVWMO may become redundant under RVTSO; the cost that the default of RVWMO imposes on Ztso implementations is the incremental overhead of fetching those fences (e.g., FENCE R,RW and FENCE RW,W) which become no-ops on that implementation. However, these fences must remain present in the code if compatibility with non-Ztso implementations is desired.

A.2. Litmus Tests

The explanations in this chapter make use of litmus tests, or small programs designed to test or highlight one particular aspect of a memory model. Litmus sample shows an example of a litmus test with two harts. As a convention for this figure and for all figures that follow in this chapter, we assume that s0-s2 are pre-set to the same value in all harts and that s0 holds the address labeled x, s1 holds y, and s2 holds z, where x, y, and z are disjoint memory locations aligned to 8 byte boundaries. All other registers and all referenced memory locations are presumed to be initialized to zero. Each figure shows the litmus test code on the left, and a visualization of one particular valid or invalid execution on the right.

Table 31. A sample litmus test and one forbidden execution (a0=1).
Hart 0 Hart 1

⋮

⋮

li t1,1

li t4,4

(a)

sw t1,0(s0)

(e)

sw t4,0(s0)

⋮

⋮

li t2,2

(b)

sw t2,0(s0)

⋮

⋮

(c)

lw a0,0(s0)

⋮

⋮

li t3,3

li t5,5

(d)

sw t3,0(s0)

(f)

sw t5,0(s0)

⋮

⋮
litmus sample

Litmus tests are used to understand the implications of the memory model in specific concrete situations. For example, in the litmus test of Litmus sample, the final value of a0 in the first hart can be either 2, 4, or 5, depending on the dynamic interleaving of the instruction stream from each hart at runtime. However, in this example, the final value of a0 in Hart 0 will never be 1 or 3; intuitively, the value 1 will no longer be visible at the time the load executes, and the value 3 will not yet be visible by the time the load executes. We analyze this test and many others below.

Table 32. A key for the litmus test diagrams drawn in this appendix
Edge Full Name (and explanation)

rf

Reads From (from each store to the loads that return a value written by that store)

co

Coherence (a total order on the stores to each address)

fr

From-Reads (from each load to co-successors of the store from which the load returned a value)

ppo

Preserved Program Order

fence

Orderings enforced by a FENCE instruction

addr

Address Dependency

ctrl

Control Dependency

data

Data Dependency

The diagram shown to the right of each litmus test shows a visual representation of the particular execution candidate being considered. These diagrams use a notation that is common in the memory model literature for constraining the set of possible global memory orders that could produce the execution in question. It is also the basis for the herd models presented in Section B.2. This notation is explained in Table 32. Of the listed relations, rf edges between harts, co edges, fr edges, and ppo edges directly constrain the global memory order (as do fence, addr, data, and some ctrl edges, via ppo). Other edges (such as intra-hart rf edges) are informative but do not constrain the global memory order.

For example, in Litmus sample, a0=1 could occur only if one of the following were true:

  • (b) appears before (a) in global memory order (and in the coherence order co). However, this violates RVWMO PPO rule ppo:→st. The co edge from (b) to (a) highlights this contradiction.

  • (a) appears before (b) in global memory order (and in the coherence order co). However, in this case, the Load Value Axiom would be violated, because (a) is not the latest matching store prior to (c) in program order. The fr edge from (c) to (b) highlights this contradiction.

Since neither of these scenarios satisfies the RVWMO axioms, the outcome a0=1 is forbidden.

Beyond what is described in this appendix, a suite of more than seven thousand litmus tests is available at github.com/litmus-tests/litmus-tests-riscv.

The litmus tests repository also provides instructions on how to run the litmus tests on RISC-V hardware and how to compare the results with the operational and axiomatic models.

In the future, we expect to adapt these memory model litmus tests for use as part of the RISC-V compliance test suite as well.

A.3. Explaining the RVWMO Rules

In this section, we provide explanation and examples for all of the RVWMO rules and axioms.

A.3.1. Preserved Program Order and Global Memory Order

Preserved program order represents the subset of program order that must be respected within the global memory order. Conceptually, events from the same hart that are ordered by preserved program order must appear in that order from the perspective of other harts and/or observers. Events from the same hart that are not ordered by preserved program order, on the other hand, may appear reordered from the perspective of other harts and/or observers.

Informally, the global memory order represents the order in which loads and stores perform. The formal memory model literature has moved away from specifications built around the concept of performing, but the idea is still useful for building up informal intuition. A load is said to have performed when its return value is determined. A store is said to have performed not when it has executed inside the pipeline, but rather only when its value has been propagated to globally visible memory. In this sense, the global memory order also represents the contribution of the coherence protocol and/or the rest of the memory system to interleave the (possibly reordered) memory accesses being issued by each hart into a single total order agreed upon by all harts.

The order in which loads perform does not always directly correspond to the relative age of the values those two loads return. In particular, a load b may perform before another load a to the same address (i.e., b may execute before a, and b may appear before a in the global memory order), but a may nevertheless return an older value than b. This discrepancy captures (among other things) the reordering effects of buffering placed between the core and memory. For example, b may have returned a value from a store in the store buffer, while a may have ignored that younger store and read an older value from memory instead. To account for this, at the time each load performs, the value it returns is determined by the load value axiom, not just strictly by determining the most recent store to the same address in the global memory order, as described below.

A.3.2. Load value axiom

Section 18.1.4.1: Each byte of each load i returns the value written to that byte by the store that is the latest in global memory order among the following stores:

  1. Stores that write that byte and that precede i in the global memory order

  2. Stores that write that byte and that precede i in program order

Preserved program order is not required to respect the ordering of a store followed by a load to an overlapping address. This complexity arises due to the ubiquity of store buffers in nearly all implementations. Informally, the load may perform (return a value) by forwarding from the store while the store is still in the store buffer, and hence before the store itself performs (writes back to globally visible memory). Any other hart will therefore observe the load as performing before the store.

Consider the Table 33. When running this program on an implementation with store buffers, it is possible to arrive at the final outcome a0=1, a1=0, a2=1, a3=0 as follows:

Table 33. A store buffer forwarding litmus test (outcome permitted)
Hart 0 Hart 1

li t1, 1

li t1, 1

(a) sw t1,0(s0)

(e) sw t1,0(s1)

(b) lw a0,0(s0)

(f) lw a2,0(s1)

(c) fence r,r

(g) fence r,r

(d) lw a1,0(s1)

(h) lw a3,0(s0)

Outcome: a0=1, a1=0, a2=1, a3=0
litmus sb fwd

  • (a) executes and enters the first hart’s private store buffer

  • (b) executes and forwards its return value 1 from (a) in the store buffer

  • (c) executes since all previous loads (i.e., (b)) have completed

  • (d) executes and reads the value 0 from memory

  • (e) executes and enters the second hart’s private store buffer

  • (f) executes and forwards its return value 1 from (e) in the store buffer

  • (g) executes since all previous loads (i.e., (f)) have completed

  • (h) executes and reads the value 0 from memory

  • (a) drains from the first hart’s store buffer to memory

  • (e) drains from the second hart’s store buffer to memory

Therefore, the memory model must be able to account for this behavior.

To put it another way, suppose the definition of preserved program order did include the following hypothetical rule: memory access a precedes memory access b in preserved program order (and hence also in the global memory order) if a precedes b in program order and a and b are accesses to the same memory location, a is a write, and b is a read. Call this "Rule X". Then we get the following:

  • (a) precedes (b): by rule X

  • (b) precedes (d): by rule 4

  • (d) precedes (e): by the load value axiom. Otherwise, if (e) preceded (d), then (d) would be required to return the value 1. (This is a perfectly legal execution; it’s just not the one in question)

  • (e) precedes (f): by rule X

  • (f) precedes (h): by rule 4

  • (h) precedes (a): by the load value axiom, as above.

The global memory order must be a total order and cannot be cyclic, because a cycle would imply that every event in the cycle happens before itself, which is impossible. Therefore, the execution proposed above would be forbidden, and hence the addition of rule X would forbid implementations with store buffer forwarding, which would clearly be undesirable.

Nevertheless, even if (b) precedes (a) and/or (f) precedes (e) in the global memory order, the only sensible possibility in this example is for (b) to return the value written by (a), and likewise for (f) and (e). This combination of circumstances is what leads to the second option in the definition of the load value axiom. Even though (b) precedes (a) in the global memory order, (a) will still be visible to (b) by virtue of sitting in the store buffer at the time (b) executes. Therefore, even if (b) precedes (a) in the global memory order, (b) should return the value written by (a) because (a) precedes (b) in program order. Likewise for (e) and (f).

Table 34. The "PPOCA" store buffer forwarding litmus test (outcome permitted)
Hart 0 Hart 1

li t1, 1

li t1, 1

(a)

sw t1,0(s0)

LOOP:

(b)

fence w,w

(d)

lw a0,0(s1)

(c)

sw t1,0(s1)

beqz a0, LOOP

(e)

sw t1,0(s2)

(f)

lw a1,0(s2)

xor a2,a1,a1

add s0,s0,a2

(g)

lw a2,0(s0)

Outcome: a0=1, a1=1, a2=0
litmus ppoca

Another test that highlights the behavior of store buffers is shown in Table 34. In this example, (d) is ordered before (e) because of the control dependency, and (f) is ordered before (g) because of the address dependency. However, (e) is not necessarily ordered before (f), even though (f) returns the value written by (e). This could correspond to the following sequence of events:

  • (e) executes speculatively and enters the second hart’s private store buffer (but does not drain to memory)

  • (f) executes speculatively and forwards its return value 1 from (e) in the store buffer

  • (g) executes speculatively and reads the value 0 from memory

  • (a) executes, enters the first hart’s private store buffer, and drains to memory

  • (b) executes and retires

  • (c) executes, enters the first hart’s private store buffer, and drains to memory

  • (d) executes and reads the value 1 from memory

  • (e), (f), and (g) commit, since the speculation turned out to be correct

  • (e) drains from the store buffer to memory

A.3.3. Atomicity axiom

Atomicity Axiom (for Aligned Atomics): If r and w are paired load and store operations generated by aligned LR and SC instructions in a hart h, s is a store to byte x, and r returns a value written by s, then s must precede w in the global memory order, and there can be no store from a hart other than h to byte x following s and preceding w in the global memory order.

The RISC-V architecture decouples the notion of atomicity from the notion of ordering. Unlike architectures such as TSO, RISC-V atomics under RVWMO do not impose any ordering requirements by default. Ordering semantics are only guaranteed by the PPO rules that otherwise apply.

RISC-V contains two types of atomics: AMOs and LR/SC pairs. These conceptually behave differently, in the following way. LR/SC behave as if the old value is brought up to the core, modified, and written back to memory, all while a reservation is held on that memory location. AMOs on the other hand conceptually behave as if they are performed directly in memory. AMOs are therefore inherently atomic, while LR/SC pairs are atomic in the slightly different sense that the memory location in question will not be modified by another hart during the time the original hart holds the reservation.

(a) lr.d a0, 0(s0) (a) lr.d a0, 0(s0) (a) lr.w a0, 0(s0) (a) lr.w a0, 0(s0)

(b) sd t1, 0(s0)

(b) sw t1, 4(s0)

(b) sw t1, 4(s0)

(b) sw t1, 4(s0)

(c) sc.d t3, t2, 0(s0)

(c) sc.d t3, t2, 0(s0)

(c) sc.w t3, t2, 0(s0)

(c) addi s0, s0, 8

(d) sc.w t3, t2, 8(s0)

Figure 4: In all four (independent) instances, the final store-conditional instruction is permitted but not guaranteed to succeed.

The atomicity axiom forbids stores from other harts from being interleaved in global memory order between an LR and the SC paired with that LR. The atomicity axiom does not forbid loads from being interleaved between the paired operations in program order or in the global memory order, nor does it forbid stores from the same hart or stores to non-overlapping locations from appearing between the paired operations in either program order or in the global memory order. For example, the SC instructions in [litmus_lrsdsc] may (but are not guaranteed to) succeed. None of those successes would violate the atomicity axiom, because the intervening non-conditional stores are from the same hart as the paired load-reserved and store-conditional instructions. This way, a memory system that tracks memory accesses at cache line granularity (and which therefore will see the four snippets of [litmus_lrsdsc] as identical) will not be forced to fail a store-conditional instruction that happens to (falsely) share another portion of the same cache line as the memory location being held by the reservation.

The atomicity axiom also technically supports cases in which the LR and SC touch different addresses and/or use different access sizes; however, use cases for such behaviors are expected to be rare in practice. Likewise, scenarios in which stores from the same hart between an LR/SC pair actually overlap the memory location(s) referenced by the LR or SC are expected to be rare compared to scenarios where the intervening store may simply fall onto the same cache line.

A.3.4. Progress axiom

Progress Axiom: No memory operation may be preceded in the global memory order by an infinite sequence of other memory operations.

The progress axiom ensures a minimal forward progress guarantee. It ensures that stores from one hart will eventually be made visible to other harts in the system in a finite amount of time, and that loads from other harts will eventually be able to read those values (or successors thereof). Without this rule, it would be legal, for example, for a spinlock to spin infinitely on a value, even with a store from another hart waiting to unlock the spinlock.

The progress axiom is intended not to impose any other notion of fairness, latency, or quality of service onto the harts in a RISC-V implementation. Any stronger notions of fairness are up to the rest of the ISA and/or up to the platform and/or device to define and implement.

The forward progress axiom will in almost all cases be naturally satisfied by any standard cache coherence protocol. Implementations with non-coherent caches may have to provide some other mechanism to ensure the eventual visibility of all stores (or successors thereof) to all harts.

A.3.5. Overlapping-Address Orderings (Rules 1-3)

Rule 1: b is a store, and a and b access overlapping memory addresses

Rule 2: a and b are loads, x is a byte read by both a and b, there is no store to x between a and b in program order, and a and b return values for x written by different memory operations

Rule 3: a is generated by an AMO or SC instruction, b is a load, and b returns a value written by a

Same-address orderings where the latter is a store are straightforward: a load or store can never be reordered with a later store to an overlapping memory location. From a microarchitecture perspective, generally speaking, it is difficult or impossible to undo a speculatively reordered store if the speculation turns out to be invalid, so such behavior is simply disallowed by the model. Same-address orderings from a store to a later load, on the other hand, do not need to be enforced. As discussed in Load value axiom, this reflects the observable behavior of implementations that forward values from buffered stores to later loads.

Same-address load-load ordering requirements are far more subtle. The basic requirement is that a younger load must not return a value that is older than a value returned by an older load in the same hart to the same address. This is often known as "CoRR" (Coherence for Read-Read pairs), or as part of a broader "coherence" or "sequential consistency per location" requirement. Some architectures in the past have relaxed same-address load-load ordering, but in hindsight this is generally considered to complicate the programming model too much, and so RVWMO requires CoRR ordering to be enforced. However, because the global memory order corresponds to the order in which loads perform rather than the ordering of the values being returned, capturing CoRR requirements in terms of the global memory order requires a bit of indirection.

Table 35. Litmus test MP+fence.w.w+fre-rfi-addr (outcome permitted)
Hart 0 Hart 1

li t1, 1

li t2, 2

(a)

sw t1,0(s0)

(d)

lw a0,0(s1)

(b)

fence w, w

(e)

sw t2,0(s1)

(c)

sw t1,0(s1)

(f)

lw a1,0(s1)

(g)

xor t3,a1,a1

(h)

add s0,s0,t3

(i)

lw a2,0(s0)

Outcome: a0=1, a1=2, a2=0
litmus mp fenceww fri rfi addr

Consider the litmus test of Table 35, which is one particular instance of the more general "fri-rfi" pattern. The term "fri-rfi" refers to the sequence (d), (e), (f): (d) "from-reads" (i.e., reads from an earlier write than) (e) which is the same hart, and (f) reads from (e) which is in the same hart.

From a microarchitectural perspective, outcome a0=1, a1=2, a2=0 is legal (as are various other less subtle outcomes). Intuitively, the following would produce the outcome in question:

  • (d) stalls (for whatever reason; perhaps it’s stalled waiting for some other preceding instruction)

  • (e) executes and enters the store buffer (but does not yet drain to memory)

  • (f) executes and forwards from (e) in the store buffer

  • (g), (h), and (i) execute

  • (a) executes and drains to memory, (b) executes, and (c) executes and drains to memory

  • (d) unstalls and executes

  • (e) drains from the store buffer to memory

This corresponds to a global memory order of (f), (i), (a), (c), (d), (e). Note that even though (f) performs before (d), the value returned by (f) is newer than the value returned by (d). Therefore, this execution is legal and does not violate the CoRR requirements.

Likewise, if two back-to-back loads return the values written by the same store, then they may also appear out-of-order in the global memory order without violating CoRR. Note that this is not the same as saying that the two loads return the same value, since two different stores may write the same value.

Table 36. Litmus test RSW (outcome permitted)
Hart 0 Hart 1

li t1, 1

(d)

lw a0,0(s1)

(a)

sw t1,0(s0)

(e)

xor t2,a0,a0

(b)

fence w, w

(f)

add s4,s2,t2

(c)

sw t1,0(s1)

(g)

lw a1,0(s4)

(h)

lw a2,0(s2)

(i)

xor t3,a2,a2

(j)

add s0,s0,t3

(k)

lw a3,0(s0)

Outcome: a0=1, a1=v, a2=v, a3=0
litmus rsw

Consider the litmus test of Table 36. The outcome a0=1, a1=v, a2=v, a3=0 (where v is some value written by another hart) can be observed by allowing (g) and (h) to be reordered. This might be done speculatively, and the speculation can be justified by the microarchitecture (e.g., by snooping for cache invalidations and finding none) because replaying (h) after (g) would return the value written by the same store anyway. Hence assuming a1 and a2 would end up with the same value written by the same store anyway, (g) and (h) can be legally reordered. The global memory order corresponding to this execution would be (h),(k),(a),(c),(d),(g).

Executions of the test in Table 36 in which a1 does not equal a2 do in fact require that (g) appears before (h) in the global memory order. Allowing (h) to appear before (g) in the global memory order would in that case result in a violation of CoRR, because then (h) would return an older value than that returned by (g). Therefore, rule 2 forbids this CoRR violation from occurring. As such, rule 2 strikes a careful balance between enforcing CoRR in all cases while simultaneously being weak enough to permit "RSW" and "fri-rfi" patterns that commonly appear in real microarchitectures.

There is one more overlapping-address rule: rule 3 simply states that a value cannot be returned from an AMO or SC to a subsequent load until the AMO or SC has (in the case of the SC, successfully) performed globally. This follows somewhat naturally from the conceptual view that both AMOs and SC instructions are meant to be performed atomically in memory. However, notably, rule 3 states that hardware may not even non-speculatively forward the value being stored by an AMOSWAP to a subsequent load, even though for AMOSWAP that store value is not actually semantically dependent on the previous value in memory, as is the case for the other AMOs. The same holds true even when forwarding from SC store values that are not semantically dependent on the value returned by the paired LR.

The three PPO rules above also apply when the memory accesses in question only overlap partially. This can occur, for example, when accesses of different sizes are used to access the same object. Note also that the base addresses of two overlapping memory operations need not necessarily be the same for two memory accesses to overlap. When misaligned memory accesses are being used, the overlapping-address PPO rules apply to each of the component memory accesses independently.

A.3.6. Fences (Rule 4)

Rule 4: There is a FENCE instruction that orders a before b

By default, the FENCE instruction ensures that all memory accesses from instructions preceding the fence in program order (the "predecessor set") appear earlier in the global memory order than memory accesses from instructions appearing after the fence in program order (the "successor set"). However, fences can optionally further restrict the predecessor set and/or the successor set to a smaller set of memory accesses in order to provide some speedup. Specifically, fences have PR, PW, SR, and SW bits which restrict the predecessor and/or successor sets. The predecessor set includes loads (resp.stores) if and only if PR (resp.PW) is set. Similarly, the successor set includes loads (resp.stores) if and only if SR (resp.SW) is set.

The FENCE encoding currently has nine non-trivial combinations of the four bits PR, PW, SR, and SW, plus one extra encoding FENCE.TSO which facilitates mapping of "acquire+release" or RVTSO semantics. The remaining seven combinations have empty predecessor and/or successor sets and hence are no-ops. Of the ten non-trivial options, only six are commonly used in practice:

  • FENCE RW,RW

  • FENCE.TSO

  • FENCE RW,W

  • FENCE R,RW

  • FENCE R,R

  • FENCE W,W

FENCE instructions using any other combination of PR, PW, SR, and SW are reserved. We strongly recommend that programmers stick to these six. Other combinations may have unknown or unexpected interactions with the memory model.

Finally, we note that since RISC-V uses a multi-copy atomic memory model, programmers can reason about fences bits in a thread-local manner. There is no complex notion of "fence cumulativity" as found in memory models that are not multi-copy atomic.

A.3.7. Explicit Synchronization (Rules 5-8)

Rule 5: a has an acquire annotation

Rule 6: b has a release annotation

Rule 7: a and b both have RCsc annotations

Rule 8: a is paired with b

An acquire operation, as would be used at the start of a critical section, requires all memory operations following the acquire in program order to also follow the acquire in the global memory order. This ensures, for example, that all loads and stores inside the critical section are up to date with respect to the synchronization variable being used to protect it. Acquire ordering can be enforced in one of two ways: with an acquire annotation, which enforces ordering with respect to just the synchronization variable itself, or with a FENCE R,RW, which enforces ordering with respect to all previous loads.

Listing 2. A spinlock with atomics
          sd           x1, (a1)     # Arbitrary unrelated store
          ld           x2, (a2)     # Arbitrary unrelated load
          li           t0, 1        # Initialize swap value.
      again:
          amoswap.w.aq t0, t0, (a0) # Attempt to acquire lock.
          bnez         t0, again    # Retry if held.
          # ...
          # Critical section.
          # ...
          amoswap.w.rl x0, x0, (a0) # Release lock by storing 0.
          sd           x3, (a3)     # Arbitrary unrelated store
          ld           x4, (a4)     # Arbitrary unrelated load

Consider Example 1. Because this example uses aq, the loads and stores in the critical section are guaranteed to appear in the global memory order after the AMOSWAP used to acquire the lock. However, assuming a0, a1, and a2 point to different memory locations, the loads and stores in the critical section may or may not appear after the "Arbitrary unrelated load" at the beginning of the example in the global memory order.

Listing 3. A spinlock with fences
          sd           x1, (a1)     # Arbitrary unrelated store
          ld           x2, (a2)     # Arbitrary unrelated load
          li           t0, 1        # Initialize swap value.
      again:
          amoswap.w    t0, t0, (a0) # Attempt to acquire lock.
          fence        r, rw        # Enforce "acquire" memory ordering
          bnez         t0, again    # Retry if held.
          # ...
          # Critical section.
          # ...
          fence        rw, w        # Enforce "release" memory ordering
          amoswap.w    x0, x0, (a0) # Release lock by storing 0.
          sd           x3, (a3)     # Arbitrary unrelated store
          ld           x4, (a4)     # Arbitrary unrelated load

Now, consider the alternative in Example 2. In this case, even though the AMOSWAP does not enforce ordering with an aq bit, the fence nevertheless enforces that the acquire AMOSWAP appears earlier in the global memory order than all loads and stores in the critical section. Note, however, that in this case, the fence also enforces additional orderings: it also requires that the "Arbitrary unrelated load" at the start of the program appears earlier in the global memory order than the loads and stores of the critical section. (This particular fence does not, however, enforce any ordering with respect to the "Arbitrary unrelated store" at the start of the snippet.) In this way, fence-enforced orderings are slightly coarser than orderings enforced by .aq.

Release orderings work exactly the same as acquire orderings, just in the opposite direction. Release semantics require all loads and stores preceding the release operation in program order to also precede the release operation in the global memory order. This ensures, for example, that memory accesses in a critical section appear before the lock-releasing store in the global memory order. Just as for acquire semantics, release semantics can be enforced using release annotations or with a FENCE RW,W operation. Using the same examples, the ordering between the loads and stores in the critical section and the "Arbitrary unrelated store" at the end of the code snippet is enforced only by the FENCE RW,W in Example 2, not by the rl in Example 1.

With RCpc annotations alone, store-release-to-load-acquire ordering is not enforced. This facilitates the porting of code written under the TSO and/or RCpc memory models. To enforce store-release-to-load-acquire ordering, the code must use store-release-RCsc and load-acquire-RCsc operations so that PPO rule 7 applies. RCpc alone is sufficient for many use cases in C/C but is insufficient for many other use cases in C/C, Java, and Linux, to name just a few examples; see Memory Porting for details.

PPO rule 8 indicates that an SC must appear after its paired LR in the global memory order. This will follow naturally from the common use of LR/SC to perform an atomic read-modify-write operation due to the inherent data dependency. However, PPO rule 8 also applies even when the value being stored does not syntactically depend on the value returned by the paired LR.

Lastly, we note that just as with fences, programmers need not worry about "cumulativity" when analyzing ordering annotations.

A.3.8. Syntactic Dependencies (Rules 9-11)

Rule 9: b has a syntactic address dependency on a

Rule 10: b has a syntactic data dependency on a

Rule 11: b is a store, and b has a syntactic control dependency on a

Dependencies from a load to a later memory operation in the same hart are respected by the RVWMO memory model. The Alpha memory model was notable for choosing not to enforce the ordering of such dependencies, but most modern hardware and software memory models consider allowing dependent instructions to be reordered too confusing and counterintuitive. Furthermore, modern code sometimes intentionally uses such dependencies as a particularly lightweight ordering enforcement mechanism.

The terms in Section 18.1.2 work as follows. Instructions are said to carry dependencies from their source register(s) to their destination register(s) whenever the value written into each destination register is a function of the source register(s). For most instructions, this means that the destination register(s) carry a dependency from all source register(s). However, there are a few notable exceptions. In the case of memory instructions, the value written into the destination register ultimately comes from the memory system rather than from the source register(s) directly, and so this breaks the chain of dependencies carried from the source register(s). In the case of unconditional jumps, the value written into the destination register comes from the current pc (which is never considered a source register by the memory model), and so likewise, JALR (the only jump with a source register) does not carry a dependency from rs1 to rd.

Listing 4. (c) has a syntactic dependency on both (a) and (b) via fflags, a destination register that both (a) and (b) implicitly accumulate into
(a) fadd f3,f1,f2
(b) fadd f6,f4,f5
(c) csrrs a0,fflags,x0

The notion of accumulating into a destination register rather than writing into it reflects the behavior of CSRs such as fflags. In particular, an accumulation into a register does not clobber any previous writes or accumulations into the same register. For example, in Listing 4, (c) has a syntactic dependency on both (a) and (b).

Like other modern memory models, the RVWMO memory model uses syntactic rather than semantic dependencies. In other words, this definition depends on the identities of the registers being accessed by different instructions, not the actual contents of those registers. This means that an address, control, or data dependency must be enforced even if the calculation could seemingly be optimized away. This choice ensures that RVWMO remains compatible with code that uses these false syntactic dependencies as a lightweight ordering mechanism.

Listing 5. A syntactic address dependency
ld a1,0(s0)
xor a2,a1,a1
add s1,s1,a2
ld a5,0(s1)

For example, there is a syntactic address dependency from the memory operation generated by the first instruction to the memory operation generated by the last instruction in Listing 5, even though a1 XOR a1 is zero and hence has no effect on the address accessed by the second load.

The benefit of using dependencies as a lightweight synchronization mechanism is that the ordering enforcement requirement is limited only to the specific two instructions in question. Other non-dependent instructions may be freely reordered by aggressive implementations. One alternative would be to use a load-acquire, but this would enforce ordering for the first load with respect to all subsequent instructions. Another would be to use a FENCE R,R, but this would include all previous and all subsequent loads, making this option more expensive.

Listing 6. A syntactic control dependency
lw x1,0(x2)
bne x1,x0,next
sw x3,0(x4)
next: sw x5,0(x6)

Control dependencies behave differently from address and data dependencies in the sense that a control dependency always extends to all instructions following the original target in program order. Consider Listing 6 the instruction at next will always execute, but the memory operation generated by that last instruction nevertheless still has a control dependency from the memory operation generated by the first instruction.

Listing 7. Another syntactic control dependency
lw x1,0(x2)
bne x1,x0,next
next: sw x3,0(x4)

Likewise, consider Listing 7. Even though both branch outcomes have the same target, there is still a control dependency from the memory operation generated by the first instruction in this snippet to the memory operation generated by the last instruction. This definition of control dependency is subtly stronger than what might be seen in other contexts (e.g., C++), but it conforms with standard definitions of control dependencies in the literature.

Notably, PPO rules 9-11 are also intentionally designed to respect dependencies that originate from the output of a successful store-conditional instruction. Typically, an SC instruction will be followed by a conditional branch checking whether the outcome was successful; this implies that there will be a control dependency from the store operation generated by the SC instruction to any memory operations following the branch. PPO rule 11 in turn implies that any subsequent store operations will appear later in the global memory order than the store operation generated by the SC. However, since control, address, and data dependencies are defined over memory operations, and since an unsuccessful SC does not generate a memory operation, no order is enforced between unsuccessful SC and its dependent instructions. Moreover, since SC is defined to carry dependencies from its source registers to rd only when the SC is successful, an unsuccessful SC has no effect on the global memory order.

Table 37. A variant of the LB litmus test (outcome forbidden)

Initial values: 0(s0)=1; 0(s2)=1

Hart 0

Hart 1

(a)

ld a0,0(s0)

(e)

ld a3,0(s2)

(b)

lr a1,0(s1)

(f)

sd a3,0(s0)

(c)

sc a2,a0,0(s1)

(d)

sd a2,0(s2)

Outcome: a0=0, a3=0
litmus lb lrsc

In addition, the choice to respect dependencies originating at store-conditional instructions ensures that certain out-of-thin-air-like behaviors will be prevented. Consider Table 37. Suppose a hypothetical implementation could occasionally make some early guarantee that a store-conditional operation will succeed. In this case, (c) could return 0 to a2 early (before actually executing), allowing the sequence (d), (e), (f), (a), and then (b) to execute, and then (c) might execute (successfully) only at that point. This would imply that (c) writes its own success value to 0(s1)! Fortunately, this situation and others like it are prevented by the fact that RVWMO respects dependencies originating at the stores generated by successful SC instructions.

We also note that syntactic dependencies between instructions only have any force when they take the form of a syntactic address, control, and/or data dependency. For example: a syntactic dependency between two F instructions via one of the accumulating CSRs in Section 18.3 does not imply that the two F instructions must be executed in order. Such a dependency would only serve to ultimately set up later a dependency from both F instructions to a later CSR instruction accessing the CSR flag in question.

A.3.9. Pipeline Dependencies (Rules 12-13)

Rule 12: b is a load, and there exists some store m between a and b in program order such that m has an address or data dependency on a, and b returns a value written by m

Rule 13: b is a store, and there exists some instruction m between a and b in program order such that m has an address dependency on a
Table 38. Because of PPO rule 12 and the data dependency from (d) to (e), (d) must also precede (f) in the global memory order (outcome forbidden)
Hart 0 Hart 1

li t1, 1

(d)

lw a0, 0(s1)

(a)

sw t1,0(s0)

(e)

sw a0, 0(s2)

(b)

fence w, w

(f)

lw a1, 0(s2)

(c)

sw t1,0(s1)

xor a2,a1,a1

add s0,s0,a2

(g)

lw a3,0(s0)

Outcome: a0=1, a3=0
litmus datarfi

PPO rules 12 and 13 reflect behaviors of almost all real processor pipeline implementations. Rule 12 states that a load cannot forward from a store until the address and data for that store are known. Consider Table 38 (f) cannot be executed until the data for (e) has been resolved, because (f) must return the value written by (e) (or by something even later in the global memory order), and the old value must not be clobbered by the writeback of (e) before (d) has had a chance to perform. Therefore, (f) will never perform before (d) has performed.

Table 39. Because of the extra store between (e) and (g), (d) no longer necessarily precedes (g) (outcome permitted)
Hart 0 Hart 1

li t1, 1

li t1, 1

(a)

sw t1,0(s0)

(d)

lw a0, 0(s1)

(b)

fence w, w

(e)

sw a0, 0(s2)

(c)

sw t1,0(s1)

(f)

sw t1, 0(s2)

(g)

lw a1, 0(s2)

xor a2,a1,a1

add s0,s0,a2

(h)

lw a3,0(s0)

Outcome: a0=1, a3=0
litmus datacoirfi

If there were another store to the same address in between (e) and (f), as in Table 40, then (f) would no longer be dependent on the data of (e) being resolved, and hence the dependency of (f) on (d), which produces the data for (e), would be broken.

Rule13 makes a similar observation to the previous rule: a store cannot be performed at memory until all previous loads that might access the same address have themselves been performed. Such a load must appear to execute before the store, but it cannot do so if the store were to overwrite the value in memory before the load had a chance to read the old value. Likewise, a store generally cannot be performed until it is known that preceding instructions will not cause an exception due to failed address resolution, and in this sense, rule 13 can be seen as somewhat of a special case of rule 11.

Table 40. Because of the address dependency from (d) to (e), (d) also precedes (f) (outcome forbidden)
Hart 0 Hart 1

li t1, 1

(a)

lw a0,0(s0)

(d)

lw a1, 0(s1)

(b)

fence rw,rw

(e)

lw a2, 0(a1)

(c)

sw s2,0(s1)

(f)

sw t1, 0(s0)

Outcome: a0=1, a1=t

litmus addrpo

Consider Table 40 (f) cannot be executed until the address for (e) is resolved, because it may turn out that the addresses match; i.e., that a1=s0. Therefore, (f) cannot be sent to memory before (d) has executed and confirmed whether the addresses do indeed overlap.

A.4. Beyond Main Memory

RVWMO does not currently attempt to formally describe how FENCE.I, SFENCE.VMA, I/O fences, and PMAs behave. All of these behaviors will be described by future formalizations. In the meantime, the behavior of FENCE.I is described in Chapter 6, the behavior of SFENCE.VMA is described in the RISC-V Instruction Set Privileged Architecture Manual, and the behavior of I/O fences and the effects of PMAs are described below.

A.4.1. Coherence and Cacheability

The RISC-V Privileged ISA defines Physical Memory Attributes (PMAs) which specify, among other things, whether portions of the address space are coherent and/or cacheable. See the RISC-V Privileged ISA Specification for the complete details. Here, we simply discuss how the various details in each PMA relate to the memory model:

  • Main memory vs.I/O, and I/O memory ordering PMAs: the memory model as defined applies to main memory regions. I/O ordering is discussed below.

  • Supported access types and atomicity PMAs: the memory model is simply applied on top of whatever primitives each region supports.

  • Cacheability PMAs: the cacheability PMAs in general do not affect the memory model. Non-cacheable regions may have more restrictive behavior than cacheable regions, but the set of allowed behaviors does not change regardless. However, some platform-specific and/or device-specific cacheability settings may differ.

  • Coherence PMAs: The memory consistency model for memory regions marked as non-coherent in PMAs is currently platform-specific and/or device-specific: the load-value axiom, the atomicity axiom, and the progress axiom all may be violated with non-coherent memory. Note however that coherent memory does not require a hardware cache coherence protocol. The RISC-V Privileged ISA Specification suggests that hardware-incoherent regions of main memory are discouraged, but the memory model is compatible with hardware coherence, software coherence, implicit coherence due to read-only memory, implicit coherence due to only one agent having access, or otherwise.

  • Idempotency PMAs: Idempotency PMAs are used to specify memory regions for which loads and/or stores may have side effects, and this in turn is used by the microarchitecture to determine, e.g., whether prefetches are legal. This distinction does not affect the memory model.

A.4.2. I/O Ordering

For I/O, the load value axiom and atomicity axiom in general do not apply, as both reads and writes might have device-specific side effects and may return values other than the value "written" by the most recent store to the same address. Nevertheless, the following preserved program order rules still generally apply for accesses to I/O memory: memory access a precedes memory access b in global memory order if a precedes b in program order and one or more of the following holds:

  1. a precedes b in preserved program order as defined in Chapter 18, with the exception that acquire and release ordering annotations apply only from one memory operation to another memory operation and from one I/O operation to another I/O operation, but not from a memory operation to an I/O nor vice versa

  2. a and b are accesses to overlapping addresses in an I/O region

  3. a and b are accesses to the same strongly ordered I/O region

  4. a and b are accesses to I/O regions, and the channel associated with the I/O region accessed by either a or b is channel 1

  5. a and b are accesses to I/O regions associated with the same channel (except for channel 0)

Note that the FENCE instruction distinguishes between main memory operations and I/O operations in its predecessor and successor sets. To enforce ordering between I/O operations and main memory operations, code must use a FENCE with PI, PO, SI, and/or SO, plus PR, PW, SR, and/or SW. For example, to enforce ordering between a write to main memory and an I/O write to a device register, a FENCE W,O or stronger is needed.

Listing 8. Ordering memory and I/O accesses
sd t0, 0(a0)
fence w,o
sd a0, 0(a1)

When a fence is in fact used, implementations must assume that the device may attempt to access memory immediately after receiving the MMIO signal, and subsequent memory accesses from that device to memory must observe the effects of all accesses ordered prior to that MMIO operation. In other words, in Listing 8, suppose 0(a0) is in main memory and 0(a1) is the address of a device register in I/O memory. If the device accesses 0(a0) upon receiving the MMIO write, then that load must conceptually appear after the first store to 0(a0) according to the rules of the RVWMO memory model. In some implementations, the only way to ensure this will be to require that the first store does in fact complete before the MMIO write is issued. Other implementations may find ways to be more aggressive, while others still may not need to do anything different at all for I/O and main memory accesses. Nevertheless, the RVWMO memory model does not distinguish between these options; it simply provides an implementation-agnostic mechanism to specify the orderings that must be enforced.

Many architectures include separate notions of "ordering" and `completion" fences, especially as it relates to I/O (as opposed to regular main memory). Ordering fences simply ensure that memory operations stay in order, while completion fences ensure that predecessor accesses have all completed before any successors are made visible. RISC-V does not explicitly distinguish between ordering and completion fences. Instead, this distinction is simply inferred from different uses of the FENCE bits.

For implementations that conform to the RISC-V Unix Platform Specification, I/O devices and DMA operations are required to access memory coherently and via strongly ordered I/O channels. Therefore, accesses to regular main memory regions that are concurrently accessed by external devices can also use the standard synchronization mechanisms. Implementations that do not conform to the Unix Platform Specification and/or in which devices do not access memory coherently will need to use mechanisms (which are currently platform-specific or device-specific) to enforce coherency.

I/O regions in the address space should be considered non-cacheable regions in the PMAs for those regions. Such regions can be considered coherent by the PMA if they are not cached by any agent.

The ordering guarantees in this section may not apply beyond a platform-specific boundary between the RISC-V cores and the device. In particular, I/O accesses sent across an external bus (e.g., PCIe) may be reordered before they reach their ultimate destination. Ordering must be enforced in such situations according to the platform-specific rules of those external devices and buses.

A.5. Code Porting and Mapping Guidelines

Table 41. Mappings from TSO operations to RISC-V operations
x86/TSO Operation RVWMO Mapping

Load

l{b|h|w|d}; fence r,rw

Store

fence rw,w; s{b|h|w|d}

Atomic RMW

amo<op>.{w|d}.aqrl OR
loop:lr.{w|d}.aq; <op>; sc.{w|d}.aqrl; bnez loop

Fence

fence rw,rw

Table 41 provides a mapping from TSO memory operations onto RISC-V memory instructions. Normal x86 loads and stores are all inherently acquire-RCpc and release-RCpc operations: TSO enforces all load-load, load-store, and store-store ordering by default. Therefore, under RVWMO, all TSO loads must be mapped onto a load followed by FENCE R,RW, and all TSO stores must be mapped onto FENCE RW,W followed by a store. TSO atomic read-modify-writes and x86 instructions using the LOCK prefix are fully ordered and can be implemented either via an AMO with both aq and rl set, or via an LR with aq set, the arithmetic operation in question, an SC with both aq and rl set, and a conditional branch checking the success condition. In the latter case, the rl annotation on the LR turns out (for non-obvious reasons) to be redundant and can be omitted.

Alternatives to Table 41 are also possible. A TSO store can be mapped onto AMOSWAP with rl set. However, since RVWMO PPO Rule 3 forbids forwarding of values from AMOs to subsequent loads, the use of AMOSWAP for stores may negatively affect performance. A TSO load can be mapped using LR with aq set: all such LR instructions will be unpaired, but that fact in and of itself does not preclude the use of LR for loads. However, again, this mapping may also negatively affect performance if it puts more pressure on the reservation mechanism than was originally intended.

Table 42. Mappings from Power operations to RISC-V operations
Power Operation RVWMO Mapping

Load

l{b|h|w|d}

Load-Reserve

lr.{w|d}

Store

s{b|h|w|d}

Store-Conditional

sc.{w|d}

lwsync

fence.tso

sync

fence rw,rw

isync

fence.i; fence r,r

Table 42 provides a mapping from Power memory operations onto RISC-V memory instructions. Power ISYNC maps on RISC-V to a FENCE.I followed by a FENCE R,R; the latter fence is needed because ISYNC is used to define a "control+control fence" dependency that is not present in RVWMO.

Table 43. Mappings from ARM operations to RISC-V operations
ARM Operation RVWMO Mapping

Load

l{b|h|w|d}

Load-Acquire

fence rw, rw; l{b|h|w|d}; fence r,rw

Load-Exclusive

lr.{w|d}

Load-Acquire-Exclusive

lr.{w|d}.aqrl

Store

s{b|h|w|d}

Store-Release

fence rw,w; s{b|h|w|d}

Store-Exclusive

sc.{w|d}

Store-Release-Exclusive

sc.{w|d}.rl

dmb

fence rw,rw

dmb.ld

fence r,rw

dmb.st

fence w,w

isb

fence.i; fence r,r

Table 43 provides a mapping from ARM memory operations onto RISC-V memory instructions. Since RISC-V does not currently have plain load and store opcodes with aq or rl annotations, ARM load-acquire and store-release operations should be mapped using fences instead. Furthermore, in order to enforce store-release-to-load-acquire ordering, there must be a FENCE RW,RW between the store-release and load-acquire; Table 43 enforces this by always placing the fence in front of each acquire operation. ARM load-exclusive and store-exclusive instructions can likewise map onto their RISC-V LR and SC equivalents, but instead of placing a FENCE RW,RW in front of an LR with aq set, we simply also set rl instead. ARM ISB maps on RISC-V to FENCE.I followed by FENCE R,R similarly to how ISYNC maps for Power.

Table 44. Mappings from Linux memory primitives to RISC-V primitives.
Linux Operation RVWMO Mapping

smp_mb()

fence rw,rw

smp_rmb()

fence r,r

smp_wmb()

fence w,w

dma_rmb()

fence r,r

dma_wmb()

fence w,w

mb()

fence iorw,iorw

rmb()

fence ri,ri

wmb()

fence wo,wo

smp_load_acquire()

l{b|h|w|d}; fence r,rw

smp_store_release()

fence.tso; s{b|h|w|d}

Linux Construct

RVWMO AMO Mapping

atomic <op> relaxed

amo <op>.{w|d}

atomic <op> acquire

amo <op>.{w|d}.aq

atomic <op> release

amo <op>.{w|d}.rl

atomic <op>

amo <op>.{w|d}.aqrl

Linux Construct

RVWMO LR/SC Mapping

atomic <op> relaxed

loop:lr.{w|d}; <op>; sc.{w|d}; bnez loop

atomic <op> acquire

loop:lr.{w|d}.aq; <op>; sc.{w|d}; bnez loop

atomic <op> release

loop:lr.{w|d}; <op>; sc.{w|d}.aqrl^*; bnez loop OR

fence.tso; loop:lr.{w|d}; <op >; sc.{w|d}^*; bnez loop

atomic <op>

loop:lr.{w|d}.aq; <op>; sc.{w|d}.aqrl; bnez loop

With regards to Table 44, other constructs (such as spinlocks) should follow accordingly. Platforms or devices with non-coherent DMA may need additional synchronization (such as cache flush or invalidate mechanisms); currently any such extra synchronization will be device-specific.

Table 44 provides a mapping of Linux memory ordering macros onto RISC-V memory instructions. The Linux fences dma_rmb() and dma_wmb() map onto FENCE R,R and FENCE W,W, respectively, since the RISC-V Unix Platform requires coherent DMA, but would be mapped onto FENCE RI,RI and FENCE WO,WO, respectively, on a platform with non-coherent DMA. Platforms with non-coherent DMA may also require a mechanism by which cache lines can be flushed and/or invalidated. Such mechanisms will be device-specific and/or standardized in a future extension to the ISA.

The Linux mappings for release operations may seem stronger than necessary, but these mappings are needed to cover some cases in which Linux requires stronger orderings than the more intuitive mappings would provide. In particular, as of the time this text is being written, Linux is actively debating whether to require load-load, load-store, and store-store orderings between accesses in one critical section and accesses in a subsequent critical section in the same hart and protected by the same synchronization object. Not all combinations of FENCE RW,W/FENCE R,RW mappings with aq/rl mappings combine to provide such orderings. There are a few ways around this problem, including:

  1. Always use FENCE RW,W/FENCE R,RW, and never use aq/rl. This suffices but is undesirable, as it defeats the purpose of the aq/rl modifiers.

  2. Always use aq/rl, and never use FENCE RW,W/FENCE R,RW. This does not currently work due to the lack of load and store opcodes with aq and rl modifiers.

  3. Strengthen the mappings of release operations such that they would enforce sufficient orderings in the presence of either type of acquire mapping. This is the currently recommended solution, and the one shown in Table 44.

RVWMO Mapping: (a) lw a0, 0(s0) (b) fence.tso // vs. fence rw,w (c) sd x0,0(s1) …​ loop: (d) amoswap.d.aq a1,t1,0(s1) bnez a1,loop (e) lw a2,0(s2)

For example, the critical section ordering rule currently being debated by the Linux community would require (a) to be ordered before (e) in Listing 9. If that will indeed be required, then it would be insufficient for (b) to map as FENCE RW,W. That said, these mappings are subject to change as the Linux Kernel Memory Model evolves.

Listing 9. Orderings between critical sections in Linux
Linux Code:
(a) int r0 = *x;
       (bc) spin_unlock(y, 0);
....
....
(d) spin_lock(y);
(e) int r1 = *z;

RVWMO Mapping:
(a) lw a0, 0(s0)
(b) fence.tso // vs. fence rw,w
(c) sd x0,0(s1)
....
loop:
(d) lr.d.aq a1,(s1)
bnez a1,loop
sc.d a1,t1,(s1)
bnez a1,loop
(e) lw a2,0(s2)

Table 45 provides a mapping of C11/C++11 atomic operations onto RISC-V memory instructions. If load and store opcodes with aq and rl modifiers are introduced, then the mappings in Table 46 will suffice. Note however that the two mappings only interoperate correctly if atomic_<op>(memory_order_seq_cst) is mapped using an LR that has both aq and rl set. Even more importantly, a Table 45 sequentially consistent store, followed by a Table 46 sequentially consistent load can be reordered unless the Table 45 mapping of stores is strengthened by either adding a second fence or mapping the store to amoswap.rl instead.

Table 45. Mappings from C/C++ primitives to RISC-V primitives.
C/C++ Construct RVWMO Mapping

Non-atomic load

l{b|h|w|d}

atomic_load(memory_order_relaxed)

l{b|h|w|d}

atomic_load(memory_order_acquire)

l{b|h|w|d}; fence r,rw

atomic_load(memory_order_seq_cst)

fence rw,rw; l{b|h|w|d}; fence r,rw

Non-atomic store

s{b|h|w|d}

atomic_store(memory_order_relaxed)

s{b|h|w|d}

atomic_store(memory_order_release)

fence rw,w; s{b|h|w|d}

atomic_store(memory_order_seq_cst)

fence rw,w; s{b|h|w|d}

atomic_thread_fence(memory_order_acquire)

fence r,rw

atomic_thread_fence(memory_order_release)

fence rw,w

atomic_thread_fence(memory_order_acq_rel)

fence.tso

atomic_thread_fence(memory_order_seq_cst)

fence rw,rw

C/C++ Construct

RVWMO AMO Mapping

atomic_<op>(memory_order_relaxed)

amo<op>.{w|d}

atomic_<op>(memory_order_acquire)

amo<op>.{w|d}.aq

atomic_<op>(memory_order_release)

amo<op>.{w|d}.rl

atomic_<op>(memory_order_acq_rel)

amo<op>.{w|d}.aqrl

atomic_<op>(memory_order_seq_cst)

amo<op>.{w|d}.aqrl

C/C++ Construct

RVWMO LR/SC Mapping

atomic_<op>(memory_order_relaxed)

loop:lr.{w|d}; <op>; sc.{w|d};

bnez loop

atomic_<op>(memory_order_acquire)

loop:lr.{w|d}.aq; <op>; sc.{w|d};

bnez loop

atomic_<op>(memory_order_release)

loop:lr.{w|d}; <op>; sc.{w|d}.rl;

bnez loop

atomic_<op>(memory_order_acq_rel)

loop:lr.{w|d}.aq; <op>; sc.{w|d}.rl;

bnez loop

atomic_<op>(memory_order_seq_cst)

loop:lr.{w|d}.aqrl; <op>;

sc.{w|d}.rl; bnez loop
Table 46. Hypothetical mappings from C/C++ primitives to RISC-V primitives, if native load-acquire and store-release opcodes are introduced.
C/C++ Construct RVWMO Mapping

Non-atomic load

l{b|h|w|d}

atomic_load(memory_order_relaxed)

l{b|h|w|d}

atomic_load(memory_order_acquire)

l{b|h|w|d}.aq

atomic_load(memory_order_seq_cst)

l{b|h|w|d}.aq

Non-atomic store

s{b|h|w|d}

atomic_store(memory_order_relaxed)

s{b|h|w|d}

atomic_store(memory_order_release)

s{b|h|w|d}.rl

atomic_store(memory_order_seq_cst)

s{b|h|w|d}.rl

atomic_thread_fence(memory_order_acquire)

fence r,rw

atomic_thread_fence(memory_order_release)

fence rw,w

atomic_thread_fence(memory_order_acq_rel)

fence.tso

atomic_thread_fence(memory_order_seq_cst)

fence rw,rw

C/C++ Construct

RVWMO AMO Mapping

atomic_<op>(memory_order_relaxed)

amo<op>.{w|d}

atomic_<op>(memory_order_acquire)

amo<op>.{w|d}.aq

atomic_<op>(memory_order_release)

amo<op>.{w|d}.rl

atomic_<op>(memory_order_acq_rel)

amo<op>.{w|d}.aqrl

atomic_<op>(memory_order_seq_cst)

amo<op>.{w|d}.aqrl

C/C++ Construct

RVWMO LR/SC Mapping

atomic_<op>(memory_order_relaxed)

lr.{w|d}; <op>; sc.{w|d}

atomic_<op>(memory_order_acquire)

lr.{w|d}.aq; <op>; sc.{w|d}

atomic_<op>(memory_order_release)

lr.{w|d}; <op>; sc.{w|d}.rl

atomic_<op>(memory_order_acq_rel)

lr.{w|d}.aq; <op>; sc.{w|d}.rl

atomic_<op>(memory_order_seq_cst)

lr.{w|d}.aq* <op>; sc.{w|d}.rl

* must be lr.{w|d}.aqrl in order to interoperate with code mapped per Table 45

Any AMO can be emulated by an LR/SC pair, but care must be taken to ensure that any PPO orderings that originate from the LR are also made to originate from the SC, and that any PPO orderings that terminate at the SC are also made to terminate at the LR. For example, the LR must also be made to respect any data dependencies that the AMO has, given that load operations do not otherwise have any notion of a data dependency. Likewise, the effect a FENCE R,R elsewhere in the same hart must also be made to apply to the SC, which would not otherwise respect that fence. The emulator may achieve this effect by simply mapping AMOs onto lr.aq; <op>; sc.aqrl, matching the mapping used elsewhere for fully ordered atomics.

These C11/C++11 mappings require the platform to provide the following Physical Memory Attributes (as defined in the RISC-V Privileged ISA) for all memory:

  • main memory

  • coherent

  • AMOArithmetic

  • RsrvEventual

Platforms with different attributes may require different mappings, or require platform-specific SW (e.g., memory-mapped I/O).

A.6. Implementation Guidelines

The RVWMO and RVTSO memory models by no means preclude microarchitectures from employing sophisticated speculation techniques or other forms of optimization in order to deliver higher performance. The models also do not impose any requirement to use any one particular cache hierarchy, nor even to use a cache coherence protocol at all. Instead, these models only specify the behaviors that can be exposed to software. Microarchitectures are free to use any pipeline design, any coherent or non-coherent cache hierarchy, any on-chip interconnect, etc., as long as the design only admits executions that satisfy the memory model rules. That said, to help people understand the actual implementations of the memory model, in this section we provide some guidelines on how architects and programmers should interpret the models' rules.

Both RVWMO and RVTSO are multi-copy atomic (or other-multi-copy-atomic): any store value that is visible to a hart other than the one that originally issued it must also be conceptually visible to all other harts in the system. In other words, harts may forward from their own previous stores before those stores have become globally visible to all harts, but no early inter-hart forwarding is permitted. Multi-copy atomicity may be enforced in a number of ways. It might hold inherently due to the physical design of the caches and store buffers, it may be enforced via a single-writer/multiple-reader cache coherence protocol, or it might hold due to some other mechanism.

Although multi-copy atomicity does impose some restrictions on the microarchitecture, it is one of the key properties keeping the memory model from becoming extremely complicated. For example, a hart may not legally forward a value from a neighbor hart’s private store buffer (unless of course it is done in such a way that no new illegal behaviors become architecturally visible). Nor may a cache coherence protocol forward a value from one hart to another until the coherence protocol has invalidated all older copies from other caches. Of course, microarchitectures may (and high-performance implementations likely will) violate these rules under the covers through speculation or other optimizations, as long as any non-compliant behaviors are not exposed to the programmer.

As a rough guideline for interpreting the PPO rules in RVWMO, we expect the following from the software perspective:

  • programmers will use PPO rules 1 and 4-8 regularly and actively.

  • expert programmers will use PPO rules 9-11 to speed up critical paths of important data structures.

  • even expert programmers will rarely if ever use PPO rules 2-3 and 12-13 directly. These are included to facilitate common microarchitectural optimizations (rule 2) and the operational formal modeling approach (rules 3 and 12-13) described in Section B.3. They also facilitate the process of porting code from other architectures that have similar rules.

We also expect the following from the hardware perspective:

  • PPO rules 1 and 3-6 reflect well-understood rules that should pose few surprises to architects.

  • PPO rule 2 reflects a natural and common hardware optimization, but one that is very subtle and hence is worth double checking carefully.

  • PPO rule 7 may not be immediately obvious to architects, but it is a standard memory model requirement

  • The load value axiom, the atomicity axiom, and PPO rules 8-13 reflect rules that most hardware implementations will enforce naturally, unless they contain extreme optimizations. Of course, implementations should make sure to double check these rules nevertheless. Hardware must also ensure that syntactic dependencies are not optimized away.

Architectures are free to implement any of the memory model rules as conservatively as they choose. For example, a hardware implementation may choose to do any or all of the following:

  • interpret all fences as if they were FENCE RW,RW (or FENCE IORW,IORW, if I/O is involved), regardless of the bits actually set

  • implement all fences with PW and SR as if they were FENCE RW,RW (or FENCE IORW,IORW, if I/O is involved), as PW with SR is the most expensive of the four possible main memory ordering components anyway

  • emulate aq and rl as described in Section A.5

  • enforcing all same-address load-load ordering, even in the presence of patterns such as fri-rfi and RSW

  • forbid any forwarding of a value from a store in the store buffer to a subsequent AMO or LR to the same address

  • forbid any forwarding of a value from an AMO or SC in the store buffer to a subsequent load to the same address

  • implement TSO on all memory accesses, and ignore any main memory fences that do not include PW and SR ordering (e.g., as Ztso implementations will do)

  • implement all atomics to be RCsc or even fully ordered, regardless of annotation

Architectures that implement RVTSO can safely do the following:

  • Ignore all fences that do not have both PW and SR (unless the fence also orders I/O)

  • Ignore all PPO rules except for rules 4 through 7, since the rest are redundant with other PPO rules under RVTSO assumptions

Other general notes:

  • Silent stores (i.e., stores that write the same value that already exists at a memory location) behave like any other store from a memory model point of view. Likewise, AMOs which do not actually change the value in memory (e.g., an AMOMAX for which the value in rs2 is smaller than the value currently in memory) are still semantically considered store operations. Microarchitectures that attempt to implement silent stores must take care to ensure that the memory model is still obeyed, particularly in cases such as RSW Section A.3.5 which tend to be incompatible with silent stores.

  • Writes may be merged (i.e., two consecutive writes to the same address may be merged) or subsumed (i.e., the earlier of two back-to-back writes to the same address may be elided) as long as the resulting behavior does not otherwise violate the memory model semantics.

The question of write subsumption can be understood from the following example:

Table 47. Write subsumption litmus test, allowed execution
Hart 0 Hart 1

li t1, 3

li t3, 2

li t2, 1

(a)

sw t1,0(s0)

(d)

lw a0,0(s1)

(b)

fence w, w

(e)

sw a0,0(s0)

(c)

sw t2,0(s1)

(f)

sw t3,0(s0)
litmus subsumption

As written, if the load (d) reads value 1, then (a) must precede (f) in the global memory order:

  • (a) precedes (c) in the global memory order because of rule 4

  • (c) precedes (d) in the global memory order because of the Load Value axiom

  • (d) precedes (e) in the global memory order because of rule 10

  • (e) precedes (f) in the global memory order because of rule 1

In other words the final value of the memory location whose address is in s0 must be 2 (the value written by the store (f)) and cannot be 3 (the value written by the store (a)).

A very aggressive microarchitecture might erroneously decide to discard (e), as (f) supersedes it, and this may in turn lead the microarchitecture to break the now-eliminated dependency between (d) and (f) (and hence also between (a) and (f)). This would violate the memory model rules, and hence it is forbidden. Write subsumption may in other cases be legal, if for example there were no data dependency between (d) and (e).

A.6.1. Possible Future Extensions

We expect that any or all of the following possible future extensions would be compatible with the RVWMO memory model:

  • "V" vector ISA extensions

  • "J" JIT extension

  • Native encodings for load and store opcodes with aq and rl set

  • Fences limited to certain addresses

  • Cache writeback/flush/invalidate/etc.instructions

A.7. Known Issues

A.7.1. Mixed-size RSW

Table 48. Mixed-size discrepancy (permitted by axiomatic models, forbidden by operational model)
Hart 0 Hart 1

li t1, 1

li t1, 1

(a)

lw a0,0(s0)

(d)

lw a1,0(s1)

(b)

fence rw,rw

(e)

amoswap.w.rl a2,t1,0(s2)

(c)

sw t1,0(s1)

(f)

ld a3,0(s2)

(g)

lw a4,4(s2)

xor a5,a4,a4

add s0,s0,a5

(h)

sw t1,0(s0)

Outcome: a0=1, a1=1, a2=0, a3=1, a4=0
Table 49. Mixed-size discrepancy (permitted by axiomatic models, forbidden by operational model)
Hart 0 Hart 1

li t1, 1

li t1, 1

(a)

lw a0,0(s0)

(d)

ld a1,0(s1)

(b)

fence rw,rw

(e)

lw a2,4(s1)

(c)

sw t1,0(s1)

xor a3,a2,a2

add s0,s0,a3

(f)

sw t1,0(s0)

Outcome: a0=1, a1=1, a2=0
Table 50. Mixed-size discrepancy (permitted by axiomatic models, forbidden by operational model)
Hart 0 Hart 1

li t1, 1

li t1, 1

(a)

lw a0,0(s0)

(d)

sw t1,4(s1)

(b)

fence rw,rw

(e)

ld a1,0(s1)

(c)

sw t1,0(s1)

(f)

lw a2,4(s1)

xor a3,a2,a2

add s0,s0,a3

(g)

sw t1,0(s0)

Outcome: a0=1, a1=0x100000001, a2=1

There is a known discrepancy between the operational and axiomatic specifications within the family of mixed-size RSW variants shown in Table 48-Table 50. To address this, we may choose to add something like the following new PPO rule: Memory operation a precedes memory operation b in preserved program order (and hence also in the global memory order) if a precedes b in program order, a and b both access regular main memory (rather than I/O regions), a is a load, b is a store, there is a load m between a and b, there is a byte x that both a and m read, there is no store between a and m that writes to x, and m precedes b in PPO. In other words, in herd syntax, we may choose to add (po-loc & rsw);ppo;[W] to PPO. Many implementations will already enforce this ordering naturally. As such, even though this rule is not official, we recommend that implementers enforce it nevertheless in order to ensure forwards compatibility with the possible future addition of this rule to RVWMO.

Appendix B: Formal Memory Model Specifications, Version 0.1

To facilitate formal analysis of RVWMO, this chapter presents a set of formalizations using different tools and modeling approaches. Any discrepancies are unintended; the expectation is that the models describe exactly the same sets of legal behaviors.

This appendix should be treated as commentary; all normative material is provided in Chapter 17 and in the rest of the main body of the ISA specification. All currently known discrepancies are listed in Section A.7. Any other discrepancies are unintentional.

B.1. Formal Axiomatic Specification in Alloy

We present a formal specification of the RVWMO memory model in Alloy (alloy.mit.edu). This model is available online at github.com/daniellustig/riscv-memory-model.

The online material also contains some litmus tests and some examples of how Alloy can be used to model check some of the mappings in Section A.5.

Listing 10. The RVWMO memory model formalized in Alloy (1/5: PPO)
// =RVWMO PPO=

// Preserved Program Order
fun ppo : Event->Event {
  // same-address ordering
  po_loc :> Store
  + rdw
  + (AMO + StoreConditional) <: rfi

  // explicit synchronization
  + ppo_fence
  + Acquire <: ^po :> MemoryEvent
  + MemoryEvent <: ^po :> Release
  + RCsc <: ^po :> RCsc
  + pair

  // syntactic dependencies
  + addrdep
  + datadep
  + ctrldep :> Store

  // pipeline dependencies
  + (addrdep+datadep).rfi
  + addrdep.^po :> Store
}

// the global memory order respects preserved program order
fact { ppo in ^gmo }
The RVWMO memory model formalized in Alloy (2/5: Axioms)
// =RVWMO axioms=

// Load Value Axiom
fun candidates[r: MemoryEvent] : set MemoryEvent {
  (r.~^gmo & Store & same_addr[r]) // writes preceding r in gmo
  + (r.^~po & Store & same_addr[r]) // writes preceding r in po
}

fun latest_among[s: set Event] : Event { s - s.~^gmo }

pred LoadValue {
  all w: Store | all r: Load |
    w->r in rf <=> w = latest_among[candidates[r]]
}

// Atomicity Axiom
pred Atomicity {
  all r: Store.~pair |            // starting from the lr,
    no x: Store & same_addr[r] |  // there is no store x to the same addr
      x not in same_hart[r]       // such that x is from a different hart,
      and x in r.~rf.^gmo         // x follows (the store r reads from) in gmo,
      and r.pair in x.^gmo        // and r follows x in gmo
}

// Progress Axiom implicit: Alloy only considers finite executions

pred RISCV_mm { LoadValue and Atomicity /* and Progress */ }
Listing 11. The RVWMO memory model formalized in Alloy (3/5: model of memory)
//Basic model of memory

sig Hart {  // hardware thread
  start : one Event
}
sig Address {}
abstract sig Event {
  po: lone Event // program order
}

abstract sig MemoryEvent extends Event {
  address: one Address,
  acquireRCpc: lone MemoryEvent,
  acquireRCsc: lone MemoryEvent,
  releaseRCpc: lone MemoryEvent,
  releaseRCsc: lone MemoryEvent,
  addrdep: set MemoryEvent,
  ctrldep: set Event,
  datadep: set MemoryEvent,
  gmo: set MemoryEvent,  // global memory order
  rf: set MemoryEvent
}
sig LoadNormal extends MemoryEvent {} // l{b|h|w|d}
sig LoadReserve extends MemoryEvent { // lr
  pair: lone StoreConditional
}
sig StoreNormal extends MemoryEvent {}       // s{b|h|w|d}
// all StoreConditionals in the model are assumed to be successful
sig StoreConditional extends MemoryEvent {}  // sc
sig AMO extends MemoryEvent {}               // amo
sig NOP extends Event {}

fun Load : Event { LoadNormal + LoadReserve + AMO }
fun Store : Event { StoreNormal + StoreConditional + AMO }

sig Fence extends Event {
  pr: lone Fence, // opcode bit
  pw: lone Fence, // opcode bit
  sr: lone Fence, // opcode bit
  sw: lone Fence  // opcode bit
}
sig FenceTSO extends Fence {}

/* Alloy encoding detail: opcode bits are either set (encoded, e.g.,
 * as f.pr in iden) or unset (f.pr not in iden).  The bits cannot be used for
 * anything else */
fact { pr + pw + sr + sw in iden }
// likewise for ordering annotations
fact { acquireRCpc + acquireRCsc + releaseRCpc + releaseRCsc in iden }
// don't try to encode FenceTSO via pr/pw/sr/sw; just use it as-is
fact { no FenceTSO.(pr + pw + sr + sw) }
Listing 12. The RVWMO memory model formalized in Alloy (4/5: Basic model rules)
// =Basic model rules=

// Ordering annotation groups
fun Acquire : MemoryEvent { MemoryEvent.acquireRCpc + MemoryEvent.acquireRCsc }
fun Release : MemoryEvent { MemoryEvent.releaseRCpc + MemoryEvent.releaseRCsc }
fun RCpc : MemoryEvent { MemoryEvent.acquireRCpc + MemoryEvent.releaseRCpc }
fun RCsc : MemoryEvent { MemoryEvent.acquireRCsc + MemoryEvent.releaseRCsc }

// There is no such thing as store-acquire or load-release, unless it's both
fact { Load & Release in Acquire }
fact { Store & Acquire in Release }

// FENCE PPO
fun FencePRSR : Fence { Fence.(pr & sr) }
fun FencePRSW : Fence { Fence.(pr & sw) }
fun FencePWSR : Fence { Fence.(pw & sr) }
fun FencePWSW : Fence { Fence.(pw & sw) }

fun ppo_fence : MemoryEvent->MemoryEvent {
    (Load  <: ^po :> FencePRSR).(^po :> Load)
  + (Load  <: ^po :> FencePRSW).(^po :> Store)
  + (Store <: ^po :> FencePWSR).(^po :> Load)
  + (Store <: ^po :> FencePWSW).(^po :> Store)
  + (Load  <: ^po :> FenceTSO) .(^po :> MemoryEvent)
  + (Store <: ^po :> FenceTSO) .(^po :> Store)
}

// auxiliary definitions
fun po_loc : Event->Event { ^po & address.~address }
fun same_hart[e: Event] : set Event { e + e.^~po + e.^po }
fun same_addr[e: Event] : set Event { e.address.~address }

// initial stores
fun NonInit : set Event { Hart.start.*po }
fun Init : set Event { Event - NonInit }
fact { Init in StoreNormal }
fact { Init->(MemoryEvent & NonInit) in ^gmo }
fact { all e: NonInit | one e.*~po.~start }  // each event is in exactly one hart
fact { all a: Address | one Init & a.~address } // one init store per address
fact { no Init <: po and no po :> Init }
Listing 13. The RVWMO memory model formalized in Alloy (5/5: Auxiliaries)
// po
fact { acyclic[po] }

// gmo
fact { total[^gmo, MemoryEvent] } // gmo is a total order over all MemoryEvents

//rf
fact { rf.~rf in iden } // each read returns the value of only one write
fact { rf in Store <: address.~address :> Load }
fun rfi : MemoryEvent->MemoryEvent { rf & (*po + *~po) }

//dep
fact { no StoreNormal <: (addrdep + ctrldep + datadep) }
fact { addrdep + ctrldep + datadep + pair in ^po }
fact { datadep in datadep :> Store }
fact { ctrldep.*po in ctrldep }
fact { no pair & (^po :> (LoadReserve + StoreConditional)).^po }
fact { StoreConditional in LoadReserve.pair } // assume all SCs succeed

// rdw
fun rdw : Event->Event {
  (Load <: po_loc :> Load)  // start with all same_address load-load pairs,
  - (~rf.rf)                // subtract pairs that read from the same store,
  - (po_loc.rfi)            // and subtract out "fri-rfi" patterns
}

// filter out redundant instances and/or visualizations
fact { no gmo & gmo.gmo } // keep the visualization uncluttered
fact { all a: Address | some a.~address }

// =Optional: opcode encoding restrictions=

// the list of blessed fences
fact { Fence in
  Fence.pr.sr
  + Fence.pw.sw
  + Fence.pr.pw.sw
  + Fence.pr.sr.sw
  + FenceTSO
  + Fence.pr.pw.sr.sw
}

pred restrict_to_current_encodings {
  no (LoadNormal + StoreNormal) & (Acquire + Release)
}

// =Alloy shortcuts=
pred acyclic[rel: Event->Event] { no iden & ^rel }
pred total[rel: Event->Event, bag: Event] {
  all disj e, e': bag | e->e' in rel + ~rel
  acyclic[rel]
}

B.2. Formal Axiomatic Specification in Herd

The tool herd takes a memory model and a litmus test as input and simulates the execution of the test on top of the memory model. Memory models are written in the domain specific language Cat. This section provides two Cat memory model of RVWMO. The first model, Listing 15, follows the global memory order, Chapter Chapter 18, definition of RVWMO, as much as is possible for a Cat model. The second model, Listing 16, is an equivalent, more efficient, partial order based RVWMO model.

The simulator herd is part of the diy tool suite — see diy.inria.fr for software and documentation. The models and more are available online at diy.inria.fr/cats7/riscv/.

Listing 14. riscv-defs.cat, a herd definition of preserved program order (1/3)
(*************)
(* Utilities *)
(*************)

(* All fence relations *)
let fence.r.r = [R];fencerel(Fence.r.r);[R]
let fence.r.w = [R];fencerel(Fence.r.w);[W]
let fence.r.rw = [R];fencerel(Fence.r.rw);[M]
let fence.w.r = [W];fencerel(Fence.w.r);[R]
let fence.w.w = [W];fencerel(Fence.w.w);[W]
let fence.w.rw = [W];fencerel(Fence.w.rw);[M]
let fence.rw.r = [M];fencerel(Fence.rw.r);[R]
let fence.rw.w = [M];fencerel(Fence.rw.w);[W]
let fence.rw.rw = [M];fencerel(Fence.rw.rw);[M]
let fence.tso =
  let f = fencerel(Fence.tso) in
  ([W];f;[W]) | ([R];f;[M])

let fence =
  fence.r.r | fence.r.w | fence.r.rw |
  fence.w.r | fence.w.w | fence.w.rw |
  fence.rw.r | fence.rw.w | fence.rw.rw |
  fence.tso

(* Same address, no W to the same address in-between *)
let po-loc-no-w = po-loc \ (po-loc?;[W];po-loc)
(* Read same write *)
let rsw = rf^-1;rf
(* Acquire, or stronger  *)
let AQ = Acq|AcqRel
(* Release or stronger *)
and RL = RelAcqRel
(* All RCsc *)
let RCsc = Acq|Rel|AcqRel
(* Amo events are both R and W, relation rmw relates paired lr/sc *)
let AMO = R & W
let StCond = range(rmw)

(*************)
(* ppo rules *)
(*************)

(* Overlapping-Address Orderings *)
let r1 = [M];po-loc;[W]
and r2 = ([R];po-loc-no-w;[R]) \ rsw
and r3 = [AMO|StCond];rfi;[R]
(* Explicit Synchronization *)
and r4 = fence
and r5 = [AQ];po;[M]
and r6 = [M];po;[RL]
and r7 = [RCsc];po;[RCsc]
and r8 = rmw
(* Syntactic Dependencies *)
and r9 = [M];addr;[M]
and r10 = [M];data;[W]
and r11 = [M];ctrl;[W]
(* Pipeline Dependencies *)
and r12 = [R];(addr|data);[W];rfi;[R]
and r13 = [R];addr;[M];po;[W]

let ppo = r1 | r2 | r3 | r4 | r5 | r6 | r7 | r8 | r9 | r10 | r11 | r12 | r13
Listing 15. riscv.cat, a herd version of the RVWMO memory model (2/3)
Total

(* Notice that herd has defined its own rf relation *)

(* Define ppo *)
include "riscv-defs.cat"

(********************************)
(* Generate global memory order *)
(********************************)

let gmo0 = (* precursor: ie build gmo as an total order that include gmo0 *)
  loc & (W\FW) * FW | # Final write after any write to the same location
  ppo |               # ppo compatible
  rfe                 # includes herd external rf (optimization)

(* Walk over all linear extensions of gmo0 *)
with  gmo from linearizations(M\IW,gmo0)

(* Add initial writes upfront -- convenient for computing rfGMO *)
let gmo = gmo | loc & IW * (M\IW)

(**********)
(* Axioms *)
(**********)

(* Compute rf according to the load value axiom, aka rfGMO *)
let WR = loc & ([W];(gmo|po);[R])
let rfGMO = WR \ (loc&([W];gmo);WR)

(* Check equality of herd rf and of rfGMO *)
empty (rf\rfGMO)|(rfGMO\rf) as RfCons

(* Atomicity axiom *)
let infloc = (gmo & loc)^-1
let inflocext = infloc & ext
let winside  = (infloc;rmw;inflocext) & (infloc;rf;rmw;inflocext) & [W]
empty winside as Atomic
Listing 16. riscv.cat, an alternative herd presentation of the RVWMO memory model (3/3)
Partial

(***************)
(* Definitions *)
(***************)

(* Define ppo *)
include "riscv-defs.cat"

(* Compute coherence relation *)
include "cos-opt.cat"

(**********)
(* Axioms *)
(**********)

(* Sc per location *)
acyclic co|rf|fr|po-loc as Coherence

(* Main model axiom *)
acyclic co|rfe|fr|ppo as Model

(* Atomicity axiom *)
empty rmw & (fre;coe) as Atomic

B.3. An Operational Memory Model

This is an alternative presentation of the RVWMO memory model in operational style. It aims to admit exactly the same extensional behavior as the axiomatic presentation: for any given program, admitting an execution if and only if the axiomatic presentation allows it.

The axiomatic presentation is defined as a predicate on complete candidate executions. In contrast, this operational presentation has an abstract microarchitectural flavor: it is expressed as a state machine, with states that are an abstract representation of hardware machine states, and with explicit out-of-order and speculative execution (but abstracting from more implementation-specific microarchitectural details such as register renaming, store buffers, cache hierarchies, cache protocols, etc.). As such, it can provide useful intuition. It can also construct executions incrementally, making it possible to interactively and randomly explore the behavior of larger examples, while the axiomatic model requires complete candidate executions over which the axioms can be checked.

The operational presentation covers mixed-size execution, with potentially overlapping memory accesses of different power-of-two byte sizes. Misaligned accesses are broken up into single-byte accesses.

The operational model, together with a fragment of the RISC-V ISA semantics (RV64I and A), are integrated into the rmem exploration tool (github.com/rems-project/rmem). rmem can explore litmus tests (see Section A.2) and small ELF binaries exhaustively, pseudorandomly and interactively. In rmem, the ISA semantics is expressed explicitly in Sail (see github.com/rems-project/sail for the Sail language, and github.com/rems-project/sail-riscv for the RISC-V ISA model), and the concurrency semantics is expressed in Lem (see github.com/rems-project/lem for the Lem language).

rmem has a command-line interface and a web-interface. The web-interface runs entirely on the client side, and is provided online together with a library of litmus tests: www.cl.cam.ac.uk/. The command-line interface is faster than the web-interface, specially in exhaustive mode.

Below is an informal introduction of the model states and transitions. The description of the formal model starts in the next subsection.

Terminology: In contrast to the axiomatic presentation, here every memory operation is either a load or a store. Hence, AMOs give rise to two distinct memory operations, a load and a store. When used in conjunction with instruction, the terms load and store refer to instructions that give rise to such memory operations. As such, both include AMO instructions. The term acquire refers to an instruction (or its memory operation) with the acquire-RCpc or acquire-RCsc annotation. The term release refers to an instruction (or its memory operation) with the release-RCpc or release-RCsc annotation.

Model states

Model states: A model state consists of a shared memory and a tuple of hart states.

Diagram

The shared memory state records all the memory store operations that have propagated so far, in the order they propagated (this can be made more efficient, but for simplicity of the presentation we keep it this way).

Each hart state consists principally of a tree of instruction instances, some of which have been finished, and some of which have not. Non-finished instruction instances can be subject to restart, e.g. if they depend on an out-of-order or speculative load that turns out to be unsound.

Conditional branch and indirect jump instructions may have multiple successors in the instruction tree. When such instruction is finished, any un-taken alternative paths are discarded.

Each instruction instance in the instruction tree has a state that includes an execution state of the intra-instruction semantics (the ISA pseudocode for this instruction). The model uses a formalization of the intra-instruction semantics in Sail. One can think of the execution state of an instruction as a representation of the pseudocode control state, pseudocode call stack, and local variable values. An instruction instance state also includes information about the instance’s memory and register footprints, its register reads and writes, its memory operations, whether it is finished, etc.

Model transitions

The model defines, for any model state, the set of allowed transitions, each of which is a single atomic step to a new abstract machine state. Execution of a single instruction will typically involve many transitions, and they may be interleaved in operational-model execution with transitions arising from other instructions. Each transition arises from a single instruction instance; it will change the state of that instance, and it may depend on or change the rest of its hart state and the shared memory state, but it does not depend on other hart states, and it will not change them. The transitions are introduced below and defined in Section B.3.5, with a precondition and a construction of the post-transition model state for each.

Transitions for all instructions:

  • Fetch instruction: This transition represents a fetch and decode of a new instruction instance, as a program order successor of a previously fetched instruction instance (or the initial fetch address).

The model assumes the instruction memory is fixed; it does not describe the behavior of self-modifying code. In particular, the Fetch instruction transition does not generate memory load operations, and the shared memory is not involved in the transition. Instead, the model depends on an external oracle that provides an opcode when given a memory location.

  • Register write: This is a write of a register value.

  • Register read: This is a read of a register value from the most recent program-order-predecessor instruction instance that writes to that register.

  • Pseudocode internal step: This covers pseudocode internal computation: arithmetic, function calls, etc.

  • Finish instruction: At this point the instruction pseudocode is done, the instruction cannot be restarted, memory accesses cannot be discarded, and all memory effects have taken place. For conditional branch and indirect jump instructions, any program order successors that were fetched from an address that is not the one that was written to the pc register are discarded, together with the sub-tree of instruction instances below them.

Transitions specific to load instructions:

  • Initiate memory load operations: At this point the memory footprint of the load instruction is provisionally known (it could change if earlier instructions are restarted) and its individual memory load operations can start being satisfied.

  • Complete load operations: At this point all the memory load operations of the instruction have been entirely satisfied and the instruction pseudocode can continue executing. A load instruction can be subject to being restarted until the transition. But, under some conditions, the model might treat a load instruction as non-restartable even before it is finished (e.g. see ).

Transitions specific to store instructions:

  • Complete store operations: At this point all the memory store operations of the instruction have been propagated to memory, and the instruction pseudocode can continue executing.

Transitions specific to sc instructions:

Transitions specific to AMO instructions:

Transitions specific to fence instructions:

The transitions labeled stem fd03c19463adc4514951b7e42ee16e38 can always be taken eagerly, as soon as their precondition is satisfied, without excluding other behavior; the stem f9d6fbc7d894d9fca4a17872ba87643c cannot. Although Fetch instruction is marked with a stem f9d6fbc7d894d9fca4a17872ba87643c, it can be taken eagerly as long as it is not taken infinitely many times.

An instance of a non-AMO load instruction, after being fetched, will typically experience the following transitions in this order:

  1. Register read

  2. Initiate memory load operations

  3. Satisfy memory load operation by forwarding from unpropagated stores and/or Satisfy memory load operation from memory (as many as needed to satisfy all the load operations of the instance)

  4. Complete load operations

  5. Register write

  6. Finish instruction

Before, between and after the transitions above, any number of Pseudocode internal step transitions may appear. In addition, a Fetch instruction transition for fetching the instruction in the next program location will be available until it is taken.

This concludes the informal description of the operational model. The following sections describe the formal operational model.

B.3.1. Intra-instruction Pseudocode Execution

The intra-instruction semantics for each instruction instance is expressed as a state machine, essentially running the instruction pseudocode. Given a pseudocode execution state, it computes the next state. Most states identify a pending memory or register operation, requested by the pseudocode, which the memory model has to do. The states are (this is a tagged union; tags in small-caps):

Load_mem(kind, address, size, load_continuation)

- memory load operation

Early_sc_fail(res_continuation)

- allow sc to fail early

Store_ea(kind, address, size, next_state)

- memory store effective address

Store_memv(mem_value, store_continuation)

- memory store value

Fence(kind, next_state)

- fence

Read_reg(reg_name, read_continuation)

- register read

Write_reg(reg_name, reg_value, next_state)

- register write

Internal(next_state)

- pseudocode internal step

Done

- end of pseudocode

Here:

  • mem_value and reg_value are lists of bytes;

  • address is an integer of XLEN bits;

for load/store, kind identifies whether it is lr/sc, acquire-RCpc/release-RCpc, acquire-RCsc/release-RCsc, acquire-release-RCsc; * for fence, kind identifies whether it is a normal or TSO, and (for normal fences) the predecessor and successor ordering bits; * reg_name identifies a register and a slice thereof (start and end bit indices); and the continuations describe how the instruction instance will continue for each value that might be provided by the surrounding memory model (the load_continuation and read_continuation take the value loaded from memory and read from the previous register write, the store_continuation takes false for an sc that failed and true in all other cases, and res_continuation takes false if the sc fails and true otherwise).

For example, given the load instruction lw x1,0(x2), an execution will typically go as follows. The initial execution state will be computed from the pseudocode for the given opcode. This can be expected to be Read_reg(x2, read_continuation). Feeding the most recently written value of register x2 (the instruction semantics will be blocked if necessary until the register value is available), say 0x4000, to read_continuation returns Load_mem(plain_load, 0x4000, 4, load_continuation). Feeding the 4-byte value loaded from memory location 0x4000, say 0x42, to load_continuation returns Write_reg(x1, 0x42, Done). Many Internal(next_state) states may appear before and between the states above.

Notice that writing to memory is split into two steps, Store_ea and Store_memv: the first one makes the memory footprint of the store provisionally known, and the second one adds the value to be stored. We ensure these are paired in the pseudocode (Store_ea followed by Store_memv), but there may be other steps between them.

It is observable that the Store_ea can occur before the value to be stored is determined. For example, for the litmus test LB+fence.r.rw+data-po to be allowed by the operational model (as it is by RVWMO), the first store in Hart 1 has to take the Store_ea step before its value is determined, so that the second store can see it is to a non-overlapping memory footprint, allowing the second store to be committed out of order without violating coherence.

The pseudocode of each instruction performs at most one store or one load, except for AMOs that perform exactly one load and one store. Those memory accesses are then split apart into the architecturally atomic units by the hart semantics (see Initiate memory load operations and Initiate memory store operation footprints below).

Informally, each bit of a register read should be satisfied from a register write by the most recent (in program order) instruction instance that can write that bit (or from the hart’s initial register state if there is no such write). Hence, it is essential to know the register write footprint of each instruction instance, which we calculate when the instruction instance is created (see the Festch instruction action of below). We ensure in the pseudocode that each instruction does at most one register write to each register bit, and also that it does not try to read a register value it just wrote.

Data-flow dependencies (address and data) in the model emerge from the fact that each register read has to wait for the appropriate register write to be executed (as described above).

B.3.2. Instruction Instance State

Each instruction instance _i has a state comprising:

  • program_loc, the memory address from which the instruction was fetched;

  • instruction_kind, identifying whether this is a load, store, AMO, fence, branch/jump or a simple instruction (this also includes a kind similar to the one described for the pseudocode execution states);

  • src_regs, the set of source _reg_name_s (including system registers), as statically determined from the pseudocode of the instruction;

  • dst_regs, the destination _reg_name_s (including system registers), as statically determined from the pseudocode of the instruction;

  • pseudocode_state (or sometimes just state for short), one of (this is a tagged union; tags in small-caps):

Plain(isa_state) - ready to make a pseudocode transition

Pending_mem_loads(load_continuation)

- requesting memory load operation(s)

Pending_mem_stores(store_continuation)

- requesting memory store operation(s)

  • reg_reads, the register reads the instance has performed, including, for each one, the register write slices it read from;

  • reg_writes, the register writes the instance has performed;

  • mem_loads, a set of memory load operations, and for each one the as-yet-unsatisfied slices (the byte indices that have not been satisfied yet), and, for the satisfied slices, the store slices (each consisting of a memory store operation and subset of its byte indices) that satisfied it.

  • mem_stores, a set of memory store operations, and for each one a flag that indicates whether it has been propagated (passed to the shared memory) or not.

  • information recording whether the instance is committed, finished, etc.

Each memory load operation includes a memory footprint (address and size). Each memory store operations includes a memory footprint, and, when available, a value.

A load instruction instance with a non-empty mem_loads, for which all the load operations are satisfied (i.e. there are no unsatisfied load slices) is said to be entirely satisfied.

Informally, an instruction instance is said to have fully determined data if the load (and sc) instructions feeding its source registers are finished. Similarly, it is said to have a fully determined memory footprint if the load (and sc) instructions feeding its memory operation address register are finished. Formally, we first define the notion of fully determined register write: a register write stem 7511f9c6a56927681a7d279e2d413cff from reg_writes of instruction instance stem 6ac91b4e7dd35551c6ea477deba5f82d is said to be fully determined if one of the following conditions hold:

  1. stem 6ac91b4e7dd35551c6ea477deba5f82d is finished; or

  2. the value written by stem 7511f9c6a56927681a7d279e2d413cff is not affected by a memory operation that stem 6ac91b4e7dd35551c6ea477deba5f82d has made (i.e. a value loaded from memory or the result of sc), and, for every register read that stem 6ac91b4e7dd35551c6ea477deba5f82d has made, that affects stem 7511f9c6a56927681a7d279e2d413cff, the register write from which stem 6ac91b4e7dd35551c6ea477deba5f82d read is fully determined (or stem 6ac91b4e7dd35551c6ea477deba5f82d read from the initial register state).

Now, an instruction instance stem 6ac91b4e7dd35551c6ea477deba5f82d is said to have fully determined data if for every register read stem 6aec6bc26afaa12bc00c3daffd500eb1 from reg_reads, the register writes that stem 6aec6bc26afaa12bc00c3daffd500eb1 reads from are fully determined. An instruction instance stem 6ac91b4e7dd35551c6ea477deba5f82d is said to have a fully determined memory footprint if for every register read stem 6aec6bc26afaa12bc00c3daffd500eb1 from reg_reads that feeds into stem 6ac91b4e7dd35551c6ea477deba5f82d’s memory operation address, the register writes that stem 6aec6bc26afaa12bc00c3daffd500eb1 reads from are fully determined.

The rmem tool records, for every register write, the set of register writes from other instructions that have been read by this instruction at the point of performing the write. By carefully arranging the pseudocode of the instructions covered by the tool we were able to make it so that this is exactly the set of register writes on which the write depends on.

B.3.3. Hart State

The model state of a single hart comprises:

  • hart_id, a unique identifier of the hart;

  • initial_register_state, the initial register value for each register;

  • initial_fetch_address, the initial instruction fetch address;

  • instruction_tree, a tree of the instruction instances that have been fetched (and not discarded), in program order.

B.3.4. Shared Memory State

The model state of the shared memory comprises a list of memory store operations, in the order they propagated to the shared memory.

When a store operation is propagated to the shared memory it is simply added to the end of the list. When a load operation is satisfied from memory, for each byte of the load operation, the most recent corresponding store slice is returned.

For most purposes, it is simpler to think of the shared memory as an array, i.e., a map from memory locations to memory store operation slices, where each memory location is mapped to a one-byte slice of the most recent memory store operation to that location. However, this abstraction is not detailed enough to properly handle the sc instruction. The RVWMO allows store operations from the same hart as the sc to intervene between the store operation of the sc and the store operations the paired lr read from. To allow such store operations to intervene, and forbid others, the array abstraction must be extended to record more information. Here, we use a list as it is very simple, but a more efficient and scalable implementations should probably use something better.

B.3.5. Transitions

Each of the paragraphs below describes a single kind of system transition. The description starts with a condition over the current system state. The transition can be taken in the current state only if the condition is satisfied. The condition is followed by an action that is applied to that state when the transition is taken, in order to generate the new system state.

B.3.5.1. Fetch instruction

A possible program-order-successor of instruction instance stem 6ac91b4e7dd35551c6ea477deba5f82d can be fetched from address loc if:

  1. it has not already been fetched, i.e., none of the immediate successors of stem 6ac91b4e7dd35551c6ea477deba5f82d in the hart’s instruction_tree are from loc; and

  2. if stem 6ac91b4e7dd35551c6ea477deba5f82d’s pseudocode has already written an address to pc, then loc must be that address, otherwise loc is:

    • for a conditional branch, the successor address or the branch target address;

    • for a (direct) jump and link instruction (jal), the target address;

    • for an indirect jump instruction (jalr), any address; and

    • for any other instruction, stem 41660742bc942374cdffc1c2016f35fb.

Action: construct a freshly initialized instruction instance stem fd12704256eca98f1c71f4fa1dde9dd8 for the instruction in the program memory at loc, with state Plain(isa_state), computed from the instruction pseudocode, including the static information available from the pseudocode such as its instruction_kind, src_regs, and dst_regs, and add stem fd12704256eca98f1c71f4fa1dde9dd8 to the hart’s instruction_tree as a successor of stem 6ac91b4e7dd35551c6ea477deba5f82d.

The possible next fetch addresses (loc) are available immediately after fetching stem 6ac91b4e7dd35551c6ea477deba5f82d and the model does not need to wait for the pseudocode to write to pc; this allows out-of-order execution, and speculation past conditional branches and jumps. For most instructions these addresses are easily obtained from the instruction pseudocode. The only exception to that is the indirect jump instruction (jalr), where the address depends on the value held in a register. In principle the mathematical model should allow speculation to arbitrary addresses here. The exhaustive search in the rmem tool handles this by running the exhaustive search multiple times with a growing set of possible next fetch addresses for each indirect jump. The initial search uses empty sets, hence there is no fetch after indirect jump instruction until the pseudocode of the instruction writes to pc, and then we use that value for fetching the next instruction. Before starting the next iteration of exhaustive search, we collect for each indirect jump (grouped by code location) the set of values it wrote to pc in all the executions in the previous search iteration, and use that as possible next fetch addresses of the instruction. This process terminates when no new fetch addresses are detected.

B.3.5.2. Initiate memory load operations

An instruction instance stem 6ac91b4e7dd35551c6ea477deba5f82d in state Plain(Load_mem(kind, address, size, load_continuation)) can always initiate the corresponding memory load operations. Action:

  1. Construct the appropriate memory load operations stem 87e245174d1300073c38508ca7876c6f:

    • if address is aligned to size then stem 87e245174d1300073c38508ca7876c6f is a single memory load operation of size bytes from address;

    • otherwise, stem 87e245174d1300073c38508ca7876c6f is a set of size memory load operations, each of one byte, from the addresses stem f82c08e4a3f84915b619aa6472e0f810.

  2. set mem_loads of stem 6ac91b4e7dd35551c6ea477deba5f82d to stem 87e245174d1300073c38508ca7876c6f; and

  3. update the state of stem 6ac91b4e7dd35551c6ea477deba5f82d to Pending_mem_loads(load_continuation).

In Section 18.1.1 it is said that misaligned memory accesses may be decomposed at any granularity. Here we decompose them to one-byte accesses as this granularity subsumes all others.

B.3.5.3. Satisfy memory load operation by forwarding from unpropagated stores

For a non-AMO load instruction instance stem 6ac91b4e7dd35551c6ea477deba5f82d in state Pending_mem_loads(load_continuation), and a memory load operation stem f1903748b6bb79515f22670686a6d90c in stem a6b8a0ea6d0d3d30d7c26a73378ca550 that has unsatisfied slices, the memory load operation can be partially or entirely satisfied by forwarding from unpropagated memory store operations by store instruction instances that are program-order-before stem 6ac91b4e7dd35551c6ea477deba5f82d if:

  1. all program-order-previous fence instructions with .sr and .pw set are finished;

  2. for every program-order-previous fence instruction, stem f9ab899994f3d644a9c2ab98a38de0c6, with .sr and .pr set, and .pw not set, if stem f9ab899994f3d644a9c2ab98a38de0c6 is not finished then all load instructions that are program-order-before stem f9ab899994f3d644a9c2ab98a38de0c6 are entirely satisfied;

  3. for every program-order-previous fence.tso instruction, stem f9ab899994f3d644a9c2ab98a38de0c6, that is not finished, all load instructions that are program-order-before stem f9ab899994f3d644a9c2ab98a38de0c6 are entirely satisfied;

  4. if stem 6ac91b4e7dd35551c6ea477deba5f82d is a load-acquire-RCsc, all program-order-previous store-releases-RCsc are finished;

  5. if stem 6ac91b4e7dd35551c6ea477deba5f82d is a load-acquire-release, all program-order-previous instructions are finished;

  6. all non-finished program-order-previous load-acquire instructions are entirely satisfied; and

  7. all program-order-previous store-acquire-release instructions are finished;

Let stem 082668fc86f03352833a0820f08bf3cc be the set of all unpropagated memory store operation slices from non-sc store instruction instances that are program-order-before stem 6ac91b4e7dd35551c6ea477deba5f82d and have already calculated the value to be stored, that overlap with the unsatisfied slices of stem f1903748b6bb79515f22670686a6d90c, and which are not superseded by intervening store operations or store operations that are read from by an intervening load. The last condition requires, for each memory store operation slice stem f85095907f5d98a9545a2a897907c982 in stem 082668fc86f03352833a0820f08bf3cc from instruction stem fd12704256eca98f1c71f4fa1dde9dd8:

  • that there is no store instruction program-order-between stem 6ac91b4e7dd35551c6ea477deba5f82d and stem fd12704256eca98f1c71f4fa1dde9dd8 with a memory store operation overlapping stem f85095907f5d98a9545a2a897907c982; and

  • that there is no load instruction program-order-between stem 6ac91b4e7dd35551c6ea477deba5f82d and stem fd12704256eca98f1c71f4fa1dde9dd8 that was satisfied from an overlapping memory store operation slice from a different hart.

Action:

  1. update stem a6b8a0ea6d0d3d30d7c26a73378ca550 to indicate that stem f1903748b6bb79515f22670686a6d90c was satisfied by stem 082668fc86f03352833a0820f08bf3cc; and

  2. restart any speculative instructions which have violated coherence as a result of this, i.e., for every non-finished instruction stem fd12704256eca98f1c71f4fa1dde9dd8 that is a program-order-successor of stem 6ac91b4e7dd35551c6ea477deba5f82d, and every memory load operation stem 8571a8a5d097d63350efc6d8da824055 of stem fd12704256eca98f1c71f4fa1dde9dd8 that was satisfied from stem 1344dde66edf219ff9e7873942a34a33, if there exists a memory store operation slice stem 80b56f226b89976549d536e6d2967ed7 in stem 1344dde66edf219ff9e7873942a34a33, and an overlapping memory store operation slice from a different memory store operation in stem 082668fc86f03352833a0820f08bf3cc, and stem 80b56f226b89976549d536e6d2967ed7 is not from an instruction that is a program-order-successor of stem 6ac91b4e7dd35551c6ea477deba5f82d, restart stem fd12704256eca98f1c71f4fa1dde9dd8 and its restart-dependents.

Where, the restart-dependents of instruction stem 7d6abfd7a8903d8827b35524704fd563 are:

  • program-order-successors of stem 7d6abfd7a8903d8827b35524704fd563 that have data-flow dependency on a register write of stem 7d6abfd7a8903d8827b35524704fd563;

  • program-order-successors of stem 7d6abfd7a8903d8827b35524704fd563 that have a memory load operation that reads from a memory store operation of stem 7d6abfd7a8903d8827b35524704fd563 (by forwarding);

  • if stem 7d6abfd7a8903d8827b35524704fd563 is a load-acquire, all the program-order-successors of stem 7d6abfd7a8903d8827b35524704fd563;

  • if stem 7d6abfd7a8903d8827b35524704fd563 is a load, for every fence, stem f9ab899994f3d644a9c2ab98a38de0c6, with .sr and .pr set, and .pw not set, that is a program-order-successor of stem 7d6abfd7a8903d8827b35524704fd563, all the load instructions that are program-order-successors of stem f9ab899994f3d644a9c2ab98a38de0c6;

  • if stem 7d6abfd7a8903d8827b35524704fd563 is a load, for every fence.tso, stem f9ab899994f3d644a9c2ab98a38de0c6, that is a program-order-successor of stem 7d6abfd7a8903d8827b35524704fd563, all the load instructions that are program-order-successors of stem f9ab899994f3d644a9c2ab98a38de0c6; and

  • (recursively) all the restart-dependents of all the instruction instances above.

Forwarding memory store operations to a memory load might satisfy only some slices of the load, leaving other slices unsatisfied.

A program-order-previous store operation that was not available when taking the transition above might make stem 082668fc86f03352833a0820f08bf3cc provisionally unsound (violating coherence) when it becomes available. That store will prevent the load from being finished (see Finish instruction), and will cause it to restart when that store operation is propagated (see Propagate store operation).

A consequence of the transition condition above is that store-release-RCsc memory store operations cannot be forwarded to load-acquire-RCsc instructions: stem 082668fc86f03352833a0820f08bf3cc does not include memory store operations from finished stores (as those must be propagated memory store operations), and the condition above requires all program-order-previous store-releases-RCsc to be finished when the load is acquire-RCsc.

B.3.5.4. Satisfy memory load operation from memory

For an instruction instance stem 6ac91b4e7dd35551c6ea477deba5f82d of a non-AMO load instruction or an AMO instruction in the context of the Satisfy, commit and propagate operations of an AMO transition, any memory load operation stem f1903748b6bb79515f22670686a6d90c in stem a6b8a0ea6d0d3d30d7c26a73378ca550 that has unsatisfied slices, can be satisfied from memory if all the conditions of <sat_by_forwarding, Satisfy memory load operation by forwarding from unpropagated stores>> are satisfied. Action: let stem 082668fc86f03352833a0820f08bf3cc be the memory store operation slices from memory covering the unsatisfied slices of stem f1903748b6bb79515f22670686a6d90c, and apply the action of Satisfy memory operation by forwarding from unpropagates stores.

Note that Satisfy memory operation by forwarding from unpropagates stores might leave some slices of the memory load operation unsatisfied, those will have to be satisfied by taking the transition again, or taking Satisfy memory load operation from memory. Satisfy memory load operation from memory, on the other hand, will always satisfy all the unsatisfied slices of the memory load operation.
B.3.5.5. Complete load operations

A load instruction instance stem 6ac91b4e7dd35551c6ea477deba5f82d in state Pending_mem_loads(load_continuation) can be completed (not to be confused with finished) if all the memory load operations stem a6b8a0ea6d0d3d30d7c26a73378ca550 are entirely satisfied (i.e. there are no unsatisfied slices). Action: update the state of stem 6ac91b4e7dd35551c6ea477deba5f82d to Plain(load_continuation(mem_value)), where mem_value is assembled from all the memory store operation slices that satisfied stem a6b8a0ea6d0d3d30d7c26a73378ca550.

B.3.5.6. Early sc fail

An sc instruction instance stem 6ac91b4e7dd35551c6ea477deba5f82d in state Plain(Early_sc_fail(res_continuation)) can always be made to fail. Action: update the state of stem 6ac91b4e7dd35551c6ea477deba5f82d to Plain(res_continuation(false)).

B.3.5.7. Paired sc

An sc instruction instance stem 6ac91b4e7dd35551c6ea477deba5f82d in state Plain(Early_sc_fail(res_continuation)) can continue its (potentially successful) execution if stem 6ac91b4e7dd35551c6ea477deba5f82d is paired with an lr. Action: update the state of stem 6ac91b4e7dd35551c6ea477deba5f82d to Plain(res_continuation(true)).

B.3.5.8. Initiate memory store operation footprints

An instruction instance stem 6ac91b4e7dd35551c6ea477deba5f82d in state Plain(Store_ea(kind, address, size, next_state)) can always announce its pending memory store operation footprint. Action:

  1. construct the appropriate memory store operations stem f85095907f5d98a9545a2a897907c982 (without the store value):

    • if address is aligned to size then stem f85095907f5d98a9545a2a897907c982 is a single memory store operation of size bytes to address;

    • otherwise, stem f85095907f5d98a9545a2a897907c982 is a set of size memory store operations, each of one-byte size, to the addresses stem f82c08e4a3f84915b619aa6472e0f810.

  2. set stem a9642b9ca23ba78275edb15d213c9b95 to stem f85095907f5d98a9545a2a897907c982; and

  3. update the state of stem 6ac91b4e7dd35551c6ea477deba5f82d to Plain(next_state).

Note that after taking the transition above the memory store operations do not yet have their values. The importance of splitting this transition from the transition below is that it allows other program-order-successor store instructions to observe the memory footprint of this instruction, and if they don’t overlap, propagate out of order as early as possible (i.e. before the data register value becomes available).

B.3.5.9. Instantiate memory store operation values

An instruction instance stem 6ac91b4e7dd35551c6ea477deba5f82d in state Plain(Store_memv(mem_value, store_continuation)) can always instantiate the values of the memory store operations stem a9642b9ca23ba78275edb15d213c9b95. Action:

  1. split mem_value between the memory store operations stem a9642b9ca23ba78275edb15d213c9b95; and

  2. update the state of stem 6ac91b4e7dd35551c6ea477deba5f82d to Pending_mem_stores(store_continuation).

B.3.5.10. Commit store instruction

An uncommitted instruction instance stem 6ac91b4e7dd35551c6ea477deba5f82d of a non-sc store instruction or an sc instruction in the context of the Commit and propagate store operation of an sc transition, in state Pending_mem_stores(store_continuation), can be committed (not to be confused with propagated) if:

  1. stem 6ac91b4e7dd35551c6ea477deba5f82d has fully determined data;

  2. all program-order-previous conditional branch and indirect jump instructions are finished;

  3. all program-order-previous fence instructions with .sw set are finished;

  4. all program-order-previous fence.tso instructions are finished;

  5. all program-order-previous load-acquire instructions are finished;

  6. all program-order-previous store-acquire-release instructions are finished;

  7. if stem 6ac91b4e7dd35551c6ea477deba5f82d is a store-release, all program-order-previous instructions are finished;

  8. all program-order-previous memory access instructions have a fully determined memory footprint;

  9. all program-order-previous store instructions, except for sc that failed, have initiated and so have non-empty mem_stores; and

  10. all program-order-previous load instructions have initiated and so have non-empty mem_loads.

Action: record that i is committed.

Notice that if condition 8 is satisfied the conditions 9 and 10 are also satisfied, or will be satisfied after taking some eager transitions. Hence, requiring them does not strengthen the model. By requiring them, we guarantee that previous memory access instructions have taken enough transitions to make their memory operations visible for the condition check of , which is the next transition the instruction will take, making that condition simpler.
B.3.5.11. Propagate store operation

For a committed instruction instance stem 6ac91b4e7dd35551c6ea477deba5f82d in state Pending_mem_stores(store_continuation), and an unpropagated memory store operation stem b406cb52bbdcb9c7799e21d181e685c1 in stem a9642b9ca23ba78275edb15d213c9b95, stem b406cb52bbdcb9c7799e21d181e685c1 can be propagated if:

  1. all memory store operations of program-order-previous store instructions that overlap with stem b406cb52bbdcb9c7799e21d181e685c1 have already propagated;

  2. all memory load operations of program-order-previous load instructions that overlap with stem b406cb52bbdcb9c7799e21d181e685c1 have already been satisfied, and (the load instructions) are non-restartable (see definition below); and

  3. all memory load operations that were satisfied by forwarding stem b406cb52bbdcb9c7799e21d181e685c1 are entirely satisfied.

Where a non-finished instruction instance stem 7d6abfd7a8903d8827b35524704fd563 is non-restartable if:

  1. there does not exist a store instruction stem 0d6d0a3f5bcc5a27d3ba5cf91524b5a4 and an unpropagated memory store operation stem b406cb52bbdcb9c7799e21d181e685c1 of stem 0d6d0a3f5bcc5a27d3ba5cf91524b5a4 such that applying the action of the Propagate store operation transition to stem b406cb52bbdcb9c7799e21d181e685c1 will result in the restart of stem 7d6abfd7a8903d8827b35524704fd563; and

  2. there does not exist a non-finished load instruction stem 65b1894f130a02fb32beeead9ca75d74 and a memory load operation stem f1903748b6bb79515f22670686a6d90c of stem 65b1894f130a02fb32beeead9ca75d74 such that applying the action of the Satisfy memory load operation by forwarding from unpropagated stores/Satisfy memory load operation from memory transition (even if stem f1903748b6bb79515f22670686a6d90c is already satisfied) to stem f1903748b6bb79515f22670686a6d90c will result in the restart of stem 7d6abfd7a8903d8827b35524704fd563.

Action:

  1. update the shared memory state with stem b406cb52bbdcb9c7799e21d181e685c1;

  2. update stem a9642b9ca23ba78275edb15d213c9b95 to indicate that stem b406cb52bbdcb9c7799e21d181e685c1 was propagated; and

  3. restart any speculative instructions which have violated coherence as a result of this, i.e., for every non-finished instruction stem fd12704256eca98f1c71f4fa1dde9dd8 program-order-after stem 6ac91b4e7dd35551c6ea477deba5f82d and every memory load operation stem 8571a8a5d097d63350efc6d8da824055 of stem fd12704256eca98f1c71f4fa1dde9dd8 that was satisfied from stem 1344dde66edf219ff9e7873942a34a33, if there exists a memory store operation slice stem 80b56f226b89976549d536e6d2967ed7 in stem 1344dde66edf219ff9e7873942a34a33 that overlaps with stem b406cb52bbdcb9c7799e21d181e685c1 and is not from stem b406cb52bbdcb9c7799e21d181e685c1, and stem 80b56f226b89976549d536e6d2967ed7 is not from a program-order-successor of stem 6ac91b4e7dd35551c6ea477deba5f82d, restart stem fd12704256eca98f1c71f4fa1dde9dd8 and its restart-dependents (see Satisfy memory load operation by forwarding from unpropagated stores).

B.3.5.12. Commit and propagate store operation of an sc

An uncommitted sc instruction instance stem 6ac91b4e7dd35551c6ea477deba5f82d, from hart stem af98b3d273f2fa50eec5140dd48d1eae, in state Pending_mem_stores(store_continuation), with a paired lr stem fd12704256eca98f1c71f4fa1dde9dd8 that has been satisfied by some store slices stem 082668fc86f03352833a0820f08bf3cc, can be committed and propagated at the same time if:

  1. stem fd12704256eca98f1c71f4fa1dde9dd8 is finished;

  2. every memory store operation that has been forwarded to stem fd12704256eca98f1c71f4fa1dde9dd8 is propagated;

  3. the conditions of Commit store instruction is satisfied;

  4. the conditions of Propagate store instruction is satisfied (notice that an sc instruction can only have one memory store operation); and

  5. for every store slice stem f85095907f5d98a9545a2a897907c982 from stem 082668fc86f03352833a0820f08bf3cc, stem f85095907f5d98a9545a2a897907c982 has not been overwritten, in the shared memory, by a store that is from a hart that is not stem af98b3d273f2fa50eec5140dd48d1eae, at any point since stem f85095907f5d98a9545a2a897907c982 was propagated to memory.

Action:

  1. apply the actions of Commit store instruction; and

  2. apply the action of Propagate store instruction.

B.3.5.13. Late sc fail

An sc instruction instance stem 6ac91b4e7dd35551c6ea477deba5f82d in state Pending_mem_stores(store_continuation), that has not propagated its memory store operation, can always be made to fail. Action:

  1. clear stem a9642b9ca23ba78275edb15d213c9b95; and

  2. update the state of stem 6ac91b4e7dd35551c6ea477deba5f82d to Plain(store_continuation(false)).

For efficiency, the rmem tool allows this transition only when it is not possible to take the Commit and propagate store operation of an sc transition. This does not affect the set of allowed final states, but when explored interactively, if the sc should fail one should use the Early sc fail transition instead of waiting for this transition.

B.3.5.14. Complete store operations

A store instruction instance stem 6ac91b4e7dd35551c6ea477deba5f82d in state Pending_mem_stores(store_continuation), for which all the memory store operations in stem a9642b9ca23ba78275edb15d213c9b95 have been propagated, can always be completed (not to be confused with finished). Action: update the state of stem 6ac91b4e7dd35551c6ea477deba5f82d to Plain(store_continuation(true)).

B.3.5.15. Satisfy, commit and propagate operations of an AMO

An AMO instruction instance stem 6ac91b4e7dd35551c6ea477deba5f82d in state Pending_mem_loads(load_continuation) can perform its memory access if it is possible to perform the following sequence of transitions with no intervening transitions:

  1. Satisfy memory load operation from memory

  2. Complere load operations

  3. Pseudocode internal step (zero or more times)

  4. Instantiate memory store operation values

  5. Commit store instruction

  6. Propagate store operation

  7. Complete store operations

and in addition, the condition of Finish instruction, with the exception of not requiring stem 6ac91b4e7dd35551c6ea477deba5f82d to be in state Plain(Done), holds after those transitions. Action: perform the above sequence of transitions (this does not include Finish instruction), one after the other, with no intervening transitions.

Notice that program-order-previous stores cannot be forwarded to the load of an AMO. This is simply because the sequence of transitions above does not include the forwarding transition. But even if it did include it, the sequence will fail when trying to do the Propagate store operation transition, as this transition requires all program-order-previous store operations to overlapping memory footprints to be propagated, and forwarding requires the store operation to be unpropagated.

In addition, the store of an AMO cannot be forwarded to a program-order-successor load. Before taking the transition above, the store operation of the AMO does not have its value and therefore cannot be forwarded; after taking the transition above the store operation is propagated and therefore cannot be forwarded.
B.3.5.16. Commit fence

A fence instruction instance stem 6ac91b4e7dd35551c6ea477deba5f82d in state Plain(Fence(kind, next_state)) can be committed if:

  1. if stem 6ac91b4e7dd35551c6ea477deba5f82d is a normal fence and it has .pr set, all program-order-previous load instructions are finished;

  2. if stem 6ac91b4e7dd35551c6ea477deba5f82d is a normal fence and it has .pw set, all program-order-previous store instructions are finished; and

  3. if stem 6ac91b4e7dd35551c6ea477deba5f82d is a fence.tso, all program-order-previous load and store instructions are finished.

Action:

  1. record that stem 6ac91b4e7dd35551c6ea477deba5f82d is committed; and

  2. update the state of stem 6ac91b4e7dd35551c6ea477deba5f82d to Plain(next_state).

B.3.5.17. Register read

An instruction instance stem 6ac91b4e7dd35551c6ea477deba5f82d in state Plain(Read_reg(reg_name, read_cont)) can do a register read of reg_name if every instruction instance that it needs to read from has already performed the expected reg_name register write.

Let read_sources include, for each bit of reg_name, the write to that bit by the most recent (in program order) instruction instance that can write to that bit, if any. If there is no such instruction, the source is the initial register value from initial_register_state. Let reg_value be the value assembled from read_sources. Action:

  1. add reg_name to stem e4c7338ef90c20ab9f1e25877c063865 with read_sources and reg_value; and

  2. update the state of stem 6ac91b4e7dd35551c6ea477deba5f82d to Plain(read_cont(reg_value)).

B.3.5.18. Register write

An instruction instance stem 6ac91b4e7dd35551c6ea477deba5f82d in state Plain(Write_reg(reg_name, reg_value, next_state)) can always do a reg_name register write. Action:

  1. add reg_name to stem 106f9ebc134bdfd3291c4c84455ac2db with stem 6dc98c92215c61f2018efb3df23af251 and reg_value; and

  2. update the state of stem 6ac91b4e7dd35551c6ea477deba5f82d to Plain(next_state).

where stem 6dc98c92215c61f2018efb3df23af251 is a pair of the set of all read_sources from stem e4c7338ef90c20ab9f1e25877c063865, and a flag that is true iff stem 6ac91b4e7dd35551c6ea477deba5f82d is a load instruction instance that has already been entirely satisfied.

B.3.5.19. Pseudocode internal step

An instruction instance stem 6ac91b4e7dd35551c6ea477deba5f82d in state Plain(Internal(next_state)) can always do that pseudocode-internal step. Action: update the state of stem 6ac91b4e7dd35551c6ea477deba5f82d to Plain(next_state).

B.3.5.20. Finish instruction

A non-finished instruction instance stem 6ac91b4e7dd35551c6ea477deba5f82d in state Plain(Done) can be finished if:

  1. if stem 6ac91b4e7dd35551c6ea477deba5f82d is a load instruction:

    1. all program-order-previous load-acquire instructions are finished;

    2. all program-order-previous fence instructions with .sr set are finished;

    3. for every program-order-previous fence.tso instruction, stem f9ab899994f3d644a9c2ab98a38de0c6, that is not finished, all load instructions that are program-order-before stem f9ab899994f3d644a9c2ab98a38de0c6 are finished; and

    4. it is guaranteed that the values read by the memory load operations of stem 6ac91b4e7dd35551c6ea477deba5f82d will not cause coherence violations, i.e., for any program-order-previous instruction instance stem fd12704256eca98f1c71f4fa1dde9dd8, let stem 6863bf3b75317b35ee75c32d6a757974 be the combined footprint of propagated memory store operations from store instructions program-order-between stem 6ac91b4e7dd35551c6ea477deba5f82d and stem fd12704256eca98f1c71f4fa1dde9dd8, and fixed memory store operations that were forwarded to stem 6ac91b4e7dd35551c6ea477deba5f82d from store instructions program-order-between stem 6ac91b4e7dd35551c6ea477deba5f82d and stem fd12704256eca98f1c71f4fa1dde9dd8 including stem fd12704256eca98f1c71f4fa1dde9dd8, and let stem f07d30fa3303b30b878e98038c45b996 be the complement of stem 6863bf3b75317b35ee75c32d6a757974 in the memory footprint of stem 6ac91b4e7dd35551c6ea477deba5f82d. If stem f07d30fa3303b30b878e98038c45b996 is not empty:

      1. stem fd12704256eca98f1c71f4fa1dde9dd8 has a fully determined memory footprint;

      2. stem fd12704256eca98f1c71f4fa1dde9dd8 has no unpropagated memory store operations that overlap with stem f07d30fa3303b30b878e98038c45b996; and

      3. if stem fd12704256eca98f1c71f4fa1dde9dd8 is a load with a memory footprint that overlaps with stem f07d30fa3303b30b878e98038c45b996, then all the memory load operations of stem fd12704256eca98f1c71f4fa1dde9dd8 that overlap with stem f07d30fa3303b30b878e98038c45b996 are satisfied and stem fd12704256eca98f1c71f4fa1dde9dd8 is non-restartable (see the Propagate store operation transition for how to determined if an instruction is non-restartable).

        Here, a memory store operation is called fixed if the store instruction has fully determined data.

  2. stem 6ac91b4e7dd35551c6ea477deba5f82d has a fully determined data; and

  3. if stem 6ac91b4e7dd35551c6ea477deba5f82d is not a fence, all program-order-previous conditional branch and indirect jump instructions are finished.

Action:

  1. if stem 6ac91b4e7dd35551c6ea477deba5f82d is a conditional branch or indirect jump instruction, discard any untaken paths of execution, i.e., remove all instruction instances that are not reachable by the branch/jump taken in instruction_tree; and

  2. record the instruction as finished, i.e., set finished to true.

B.3.6. Limitations

  • The model covers user-level RV64I and RV64A. In particular, it does not support the misaligned atomicity granule PMA or the total store ordering extension "Ztso". It should be trivial to adapt the model to RV32I/A and to the G, Q and C extensions, but we have never tried it. This will involve, mostly, writing Sail code for the instructions, with minimal, if any, changes to the concurrency model.

  • The model covers only normal memory accesses (it does not handle I/O accesses).

  • The model does not cover TLB-related effects.

  • The model assumes the instruction memory is fixed. In particular, the Fetch instruction transition does not generate memory load operations, and the shared memory is not involved in the transition. Instead, the model depends on an external oracle that provides an opcode when given a memory location.

  • The model does not cover exceptions, traps and interrupts.

Appendix C: Vector Assembly Code Examples

CV32A65X: This appendix is not applicable as vector extension is not supported.

Appendix D: Calling Convention for Vector State (Not authoritative - Placeholder Only)

CV32A65X: This appendix is not applicable as vector extension is not supported.

Index

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