The LLVM Target-Independent Code Generator¶
Warning
This is a work in progress.
Introduction¶
The LLVM target-independent code generator is a framework that provides a suite of reusable components for translating the LLVM internal representation to the machine code for a specified target—either in assembly form (suitable for a static compiler) or in binary machine code format (usable for a JIT compiler). The LLVM target-independent code generator consists of six main components:
Abstract target description interfaces which capture important properties about various aspects of the machine, independently of how they will be used. These interfaces are defined in
include/llvm/Target/.Classes used to represent the code being generated for a target. These classes are intended to be abstract enough to represent the machine code for any target machine. These classes are defined in
include/llvm/CodeGen/. At this level, concepts like “constant pool entries” and “jump tables” are explicitly exposed.Classes and algorithms used to represent code at the object file level, the MC Layer. These classes represent assembly level constructs like labels, sections, and instructions. At this level, concepts like “constant pool entries” and “jump tables” don’t exist.
Target-independent algorithms used to implement various phases of native code generation (register allocation, scheduling, stack frame representation, etc). This code lives in
lib/CodeGen/.Implementations of the abstract target description interfaces for particular targets. These machine descriptions make use of the components provided by LLVM, and can optionally provide custom target-specific passes, to build complete code generators for a specific target. Target descriptions live in
lib/Target/.The target-independent JIT components. The LLVM JIT is completely target independent (it uses the
TargetJITInfostructure to interface for target-specific issues. The code for the target-independent JIT lives inlib/ExecutionEngine/JIT.
Depending on which part of the code generator you are interested in working on, different pieces of this will be useful to you. In any case, you should be familiar with the target description and machine code representation classes. If you want to add a backend for a new target, you will need to implement the target description classes for your new target and understand the LLVM code representation. If you are interested in implementing a new code generation algorithm, it should only depend on the target-description and machine code representation classes, ensuring that it is portable.
Required components in the code generator¶
The two pieces of the LLVM code generator are the high-level interface to the code generator and the set of reusable components that can be used to build target-specific backends. The two most important interfaces ( TargetMachine and DataLayout ) are the only ones that are required to be defined for a backend to fit into the LLVM system, but the others must be defined if the reusable code generator components are going to be used.
This design has two important implications. The first is that LLVM can support completely non-traditional code generation targets. For example, the C backend does not require register allocation, instruction selection, or any of the other standard components provided by the system. As such, it only implements these two interfaces, and does its own thing. Note that C backend was removed from the trunk since LLVM 3.1 release. Another example of a code generator like this is a (purely hypothetical) backend that converts LLVM to the GCC RTL form and uses GCC to emit machine code for a target.
This design also implies that it is possible to design and implement radically different code generators in the LLVM system that do not make use of any of the built-in components. Doing so is not recommended at all, but could be required for radically different targets that do not fit into the LLVM machine description model: FPGAs for example.
The high-level design of the code generator¶
The LLVM target-independent code generator is designed to support efficient and quality code generation for standard register-based microprocessors. Code generation in this model is divided into the following stages:
Instruction Selection — This phase determines an efficient way to express the input LLVM code in the target instruction set. This stage produces the initial code for the program in the target instruction set, then makes use of virtual registers in SSA form and physical registers that represent any required register assignments due to target constraints or calling conventions. This step turns the LLVM code into a DAG of target instructions.
Scheduling and Formation — This phase takes the DAG of target instructions produced by the instruction selection phase, determines an ordering of the instructions, then emits the instructions as MachineInstrs with that ordering. Note that we describe this in the instruction selection section because it operates on a SelectionDAG.
SSA-based Machine Code Optimizations — This optional stage consists of a series of machine-code optimizations that operate on the SSA-form produced by the instruction selector. Optimizations like modulo-scheduling or peephole optimization work here.
Register Allocation — The target code is transformed from an infinite virtual register file in SSA form to the concrete register file used by the target. This phase introduces spill code and eliminates all virtual register references from the program.
Prolog/Epilog Code Insertion — Once the machine code has been generated for the function and the amount of stack space required is known (used for LLVM alloca’s and spill slots), the prolog and epilog code for the function can be inserted and “abstract stack location references” can be eliminated. This stage is responsible for implementing optimizations like frame-pointer elimination and stack packing.
Late Machine Code Optimizations — Optimizations that operate on “final” machine code can go here, such as spill code scheduling and peephole optimizations.
Code Emission — The final stage actually puts out the code for the current function, either in the target assembler format or in machine code.
The code generator is based on the assumption that the instruction selector will use an optimal pattern matching selector to create high-quality sequences of native instructions. Alternative code generator designs based on pattern expansion and aggressive iterative peephole optimization are much slower. This design permits efficient compilation (important for JIT environments) and aggressive optimization (used when generating code offline) by allowing components of varying levels of sophistication to be used for any step of compilation.
In addition to these stages, target implementations can insert arbitrary target-specific passes into the flow. For example, the X86 target uses a special pass to handle the 80x87 floating point stack architecture. Other targets with unusual requirements can be supported with custom passes as needed.
Using TableGen for target description¶
The target description classes require a detailed description of the target
architecture. These target descriptions often have a large amount of common
information (e.g., an add instruction is almost identical to a sub
instruction). In order to allow the maximum amount of commonality to be
factored out, the LLVM code generator uses the
TableGen Overview tool to describe big chunks of the
target machine, which allows the use of domain-specific and target-specific
abstractions to reduce the amount of repetition.
As LLVM continues to be developed and refined, we plan to move more and more of
the target description to the .td form. Doing so gives us a number of
advantages. The most important is that it makes it easier to port LLVM because
it reduces the amount of C++ code that has to be written, and the surface area
of the code generator that needs to be understood before someone can get
something working. Second, it makes it easier to change things. In particular,
if tables and other things are all emitted by tblgen, we only need a change
in one place (tblgen) to update all of the targets to a new interface.
Target description classes¶
The LLVM target description classes (located in the include/llvm/Target
directory) provide an abstract description of the target machine independent of
any particular client. These classes are designed to capture the abstract
properties of the target (such as the instructions and registers it has), and do
not incorporate any particular pieces of code generation algorithms.
All of the target description classes (except the DataLayout class) are designed to be subclassed by the concrete target implementation, and have virtual methods implemented. To get to these implementations, the TargetMachine class provides accessors that should be implemented by the target.
The TargetMachine class¶
The TargetMachine class provides virtual methods that are used to access the
target-specific implementations of the various target description classes via
the get*Info methods (getInstrInfo, getRegisterInfo,
getFrameInfo, etc.). This class is designed to be specialized by a concrete
target implementation (e.g., X86TargetMachine) which implements the various
virtual methods. The only required target description class is the
DataLayout class, but if the code
generator components are to be used, the other interfaces should be implemented
as well.
The DataLayout class¶
The DataLayout class is the only required target description class, and it
is the only class that is not extensible (you cannot derive a new class from
it). DataLayout specifies information about how the target lays out memory
for structures, the alignment requirements for various data types, the size of
pointers in the target, and whether the target is little-endian or
big-endian.
The TargetLowering class¶
The TargetLowering class is used by SelectionDAG based instruction selectors
primarily to describe how LLVM code should be lowered to SelectionDAG
operations. Among other things, this class indicates:
an initial register class to use for various
ValueTypes,which operations are natively supported by the target machine,
the return type of
setccoperations,the type to use for shift amounts, and
various high-level characteristics, like whether it is profitable to turn division by a constant into a multiplication sequence.
The TargetRegisterInfo class¶
The TargetRegisterInfo class is used to describe the register file of the
target and any interactions between the registers.
Registers are represented in the code generator by unsigned integers. Physical
registers (those that actually exist in the target description) are unique
small numbers, and virtual registers are generally large. Note that
register #0 is reserved as a flag value.
Each register in the processor description has an associated
TargetRegisterDesc entry, which provides a textual name for the register
(used for assembly output and debugging dumps) and a set of aliases (used to
indicate whether one register overlaps with another).
In addition to the per-register description, the TargetRegisterInfo class
exposes a set of processor specific register classes (instances of the
TargetRegisterClass class). Each register class contains sets of registers
that have the same properties (for example, they are all 32-bit integer
registers). Each SSA virtual register created by the instruction selector has
an associated register class. When the register allocator runs, it replaces
virtual registers with a physical register in the set.
The target-specific implementations of these classes is auto-generated from a TableGen Overview description of the register file.
The TargetInstrInfo class¶
The TargetInstrInfo class is used to describe the machine instructions
supported by the target. Descriptions define things like the mnemonic for
the opcode, the number of operands, the list of implicit register uses and defs,
whether the instruction has certain target-independent properties (accesses
memory, is commutable, etc), and holds any target-specific flags.
The TargetFrameLowering class¶
The TargetFrameLowering class is used to provide information about the stack
frame layout of the target. It holds the direction of stack growth, the known
stack alignment on entry to each function, and the offset to the local area.
The offset to the local area is the offset from the stack pointer on function
entry to the first location where function data (local variables, spill
locations) can be stored.
The TargetSubtarget class¶
The TargetSubtarget class is used to provide information about the specific
chip set being targeted. A sub-target informs code generation of which
instructions are supported, instruction latencies and instruction execution
itinerary; i.e., which processing units are used, in what order, and for how
long.
The TargetJITInfo class¶
The TargetJITInfo class exposes an abstract interface used by the
Just-In-Time code generator to perform target-specific activities, such as
emitting stubs. If a TargetMachine supports JIT code generation, it should
provide one of these objects through the getJITInfo method.
Machine code description classes¶
At the high-level, LLVM code is translated to a machine specific representation
formed out of MachineFunction ,
MachineBasicBlock , and
MachineInstr instances (defined in
include/llvm/CodeGen). This representation is completely target agnostic,
representing instructions in their most abstract form: an opcode and a series of
operands. This representation is designed to support both an SSA representation
for machine code, as well as a register allocated, non-SSA form.
The MachineInstr class¶
Target machine instructions are represented as instances of the MachineInstr
class. This class is an extremely abstract way of representing machine
instructions. In particular, it only keeps track of an opcode number and a set
of operands.
The opcode number is a simple unsigned integer that only has meaning to a
specific backend. All of the instructions for a target should be defined in the
*InstrInfo.td file for the target. The opcode enum values are auto-generated
from this description. The MachineInstr class does not have any information
about how to interpret the instruction (i.e., what the semantics of the
instruction are); for that you must refer to the
TargetInstrInfo class.
The operands of a machine instruction can be of several different types: a register reference, a constant integer, a basic block reference, etc. In addition, a machine operand should be marked as a def or a use of the value (though only registers are allowed to be defs).
By convention, the LLVM code generator orders instruction operands so that all
register definitions come before the register uses, even on architectures that
are normally printed in other orders. For example, the SPARC add instruction:
“add %i1, %i2, %i3” adds the “%i1”, and “%i2” registers and stores the
result into the “%i3” register. In the LLVM code generator, the operands should
be stored as “%i3, %i1, %i2”: with the destination first.
Keeping destination (definition) operands at the beginning of the operand list has several advantages. In particular, the debugging printer will print the instruction like this:
%r3 = add %i1, %i2
Also if the first operand is a def, it is easier to create instructions whose only def is the first operand.
Using the MachineInstrBuilder.h functions¶
Machine instructions are created by using the BuildMI functions, located in
the include/llvm/CodeGen/MachineInstrBuilder.h file. The BuildMI
functions make it easy to build arbitrary machine instructions. Usage of the
BuildMI functions look like this:
// Create a 'DestReg = mov 42' (rendered in X86 assembly as 'mov DestReg, 42')
// instruction and insert it at the end of the given MachineBasicBlock.
const TargetInstrInfo &TII = ...
MachineBasicBlock &MBB = ...
DebugLoc DL;
MachineInstr *MI = BuildMI(MBB, DL, TII.get(X86::MOV32ri), DestReg).addImm(42);
// Create the same instr, but insert it before a specified iterator point.
MachineBasicBlock::iterator MBBI = ...
BuildMI(MBB, MBBI, DL, TII.get(X86::MOV32ri), DestReg).addImm(42);
// Create a 'cmp Reg, 0' instruction, no destination reg.
MI = BuildMI(MBB, DL, TII.get(X86::CMP32ri8)).addReg(Reg).addImm(42);
// Create an 'sahf' instruction which takes no operands and stores nothing.
MI = BuildMI(MBB, DL, TII.get(X86::SAHF));
// Create a self looping branch instruction.
BuildMI(MBB, DL, TII.get(X86::JNE)).addMBB(&MBB);
If you need to add a definition operand (other than the optional destination register), you must explicitly mark it as such:
MI.addReg(Reg, RegState::Define);
Fixed (preassigned) registers¶
One important issue that the code generator needs to be aware of is the presence
of fixed registers. In particular, there are often places in the instruction
stream where the register allocator must arrange for a particular value to be
in a particular register. This can occur due to limitations of the instruction
set (e.g., the X86 can only do a 32-bit divide with the EAX/EDX
registers), or external factors like calling conventions. In any case, the
instruction selector should emit code that copies a virtual register into or out
of a physical register when needed.
For example, consider this simple LLVM example:
define i32 @test(i32 %X, i32 %Y) {
%Z = sdiv i32 %X, %Y
ret i32 %Z
}
The X86 instruction selector might produce this machine code for the div and
ret:
;; Start of div
%EAX = mov %reg1024 ;; Copy X (in reg1024) into EAX
%reg1027 = sar %reg1024, 31
%EDX = mov %reg1027 ;; Sign extend X into EDX
idiv %reg1025 ;; Divide by Y (in reg1025)
%reg1026 = mov %EAX ;; Read the result (Z) out of EAX
;; Start of ret
%EAX = mov %reg1026 ;; 32-bit return value goes in EAX
ret
By the end of code generation, the register allocator would coalesce the registers and delete the resultant identity moves producing the following code:
;; X is in EAX, Y is in ECX
mov %EAX, %EDX
sar %EDX, 31
idiv %ECX
ret
This approach is extremely general (if it can handle the X86 architecture, it can handle anything!) and allows all of the target specific knowledge about the instruction stream to be isolated in the instruction selector. Note that physical registers should have a short lifetime for good code generation, and all physical registers are assumed dead on entry to and exit from basic blocks (before register allocation). Thus, if you need a value to be live across basic block boundaries, it must live in a virtual register.
Call-clobbered registers¶
Some machine instructions, like calls, clobber a large number of physical
registers. Rather than adding <def,dead> operands for all of them, it is
possible to use an MO_RegisterMask operand instead. The register mask
operand holds a bit mask of preserved registers, and everything else is
considered to be clobbered by the instruction.
Machine code in SSA form¶
MachineInstr’s are initially selected in SSA-form, and are maintained in
SSA-form until register allocation happens. For the most part, this is
trivially simple since LLVM is already in SSA form; LLVM PHI nodes become
machine code PHI nodes, and virtual registers are only allowed to have a single
definition.
After register allocation, machine code is no longer in SSA-form because there are no virtual registers left in the code.
The MachineBasicBlock class¶
The MachineBasicBlock class contains a list of machine instructions
( MachineInstr instances). It roughly
corresponds to the LLVM code input to the instruction selector, but there can be
a one-to-many mapping (i.e. one LLVM basic block can map to multiple machine
basic blocks). The MachineBasicBlock class has a “getBasicBlock” method,
which returns the LLVM basic block that it comes from.
The MachineFunction class¶
The MachineFunction class contains a list of machine basic blocks
( MachineBasicBlock instances). It
corresponds one-to-one with the LLVM function input to the instruction selector.
In addition to a list of basic blocks, the MachineFunction contains a
MachineConstantPool, a MachineFrameInfo, a MachineFunctionInfo, and
a MachineRegisterInfo. See include/llvm/CodeGen/MachineFunction.h for
more information.
MachineInstr Bundles¶
LLVM code generator can model sequences of instructions as MachineInstr bundles. A MI bundle can model a VLIW group / pack which contains an arbitrary number of parallel instructions. It can also be used to model a sequential list of instructions (potentially with data dependencies) that cannot be legally separated (e.g. ARM Thumb2 IT blocks).
Conceptually a MI bundle is a MI with a number of other MIs nested within:
--------------
| Bundle | ---------
-------------- \
| ----------------
| | MI |
| ----------------
| |
| ----------------
| | MI |
| ----------------
| |
| ----------------
| | MI |
| ----------------
|
--------------
| Bundle | --------
-------------- \
| ----------------
| | MI |
| ----------------
| |
| ----------------
| | MI |
| ----------------
| |
| ...
|
--------------
| Bundle | --------
-------------- \
|
...
MI bundle support does not change the physical representations of MachineBasicBlock and MachineInstr. All the MIs (including top level and nested ones) are stored as sequential list of MIs. The “bundled” MIs are marked with the ‘InsideBundle’ flag. A top level MI with the special BUNDLE opcode is used to represent the start of a bundle. It’s legal to mix BUNDLE MIs with individual MIs that are not inside bundles nor represent bundles.
MachineInstr passes should operate on a MI bundle as a single unit. Member methods have been taught to correctly handle bundles and MIs inside bundles. The MachineBasicBlock iterator has been modified to skip over bundled MIs to enforce the bundle-as-a-single-unit concept. An alternative iterator instr_iterator has been added to MachineBasicBlock to allow passes to iterate over all of the MIs in a MachineBasicBlock, including those which are nested inside bundles. The top level BUNDLE instruction must have the correct set of register MachineOperand’s that represent the cumulative inputs and outputs of the bundled MIs.
Packing / bundling of MachineInstrs for VLIW architectures should generally be done as part of the register allocation super-pass. More specifically, the pass which determines what MIs should be bundled together should be done after code generator exits SSA form (i.e. after two-address pass, PHI elimination, and copy coalescing). Such bundles should be finalized (i.e. adding BUNDLE MIs and input and output register MachineOperands) after virtual registers have been rewritten into physical registers. This eliminates the need to add virtual register operands to BUNDLE instructions which would effectively double the virtual register def and use lists. Bundles may use virtual registers and be formed in SSA form, but may not be appropriate for all use cases.
The “MC” Layer¶
The MC Layer is used to represent and process code at the raw machine code level, devoid of “high level” information like “constant pools”, “jump tables”, “global variables” or anything like that. At this level, LLVM handles things like label names, machine instructions, and sections in the object file. The code in this layer is used for a number of important purposes: the tail end of the code generator uses it to write a .s or .o file, and it is also used by the llvm-mc tool to implement standalone machine code assemblers and disassemblers.
This section describes some of the important classes. There are also a number of important subsystems that interact at this layer, they are described later in this manual.
The MCStreamer API¶
MCStreamer is best thought of as an assembler API. It is an abstract API which is implemented in different ways (e.g. to output a .s file, output an ELF .o file, etc) but whose API correspond directly to what you see in a .s file. MCStreamer has one method per directive, such as EmitLabel, EmitSymbolAttribute, switchSection, emitValue (for .byte, .word), etc, which directly correspond to assembly level directives. It also has an EmitInstruction method, which is used to output an MCInst to the streamer.
This API is most important for two clients: the llvm-mc stand-alone assembler is effectively a parser that parses a line, then invokes a method on MCStreamer. In the code generator, the Code Emission phase of the code generator lowers higher level LLVM IR and Machine* constructs down to the MC layer, emitting directives through MCStreamer.
On the implementation side of MCStreamer, there are two major implementations:
one for writing out a .s file (MCAsmStreamer), and one for writing out a .o
file (MCObjectStreamer). MCAsmStreamer is a straightforward implementation
that prints out a directive for each method (e.g. EmitValue -> .byte), but
MCObjectStreamer implements a full assembler.
For target specific directives, the MCStreamer has a MCTargetStreamer instance.
Each target that needs it defines a class that inherits from it and is a lot
like MCStreamer itself: It has one method per directive and two classes that
inherit from it, a target object streamer and a target asm streamer. The target
asm streamer just prints it (emitFnStart -> .fnstart), and the object
streamer implement the assembler logic for it.
To make llvm use these classes, the target initialization must call TargetRegistry::RegisterAsmStreamer and TargetRegistry::RegisterMCObjectStreamer passing callbacks that allocate the corresponding target streamer and pass it to createAsmStreamer or to the appropriate object streamer constructor.
The MCContext class¶
The MCContext class is the owner of a variety of uniqued data structures at the MC layer, including symbols, sections, etc. As such, this is the class that you interact with to create symbols and sections. This class can not be subclassed.
The MCSymbol class¶
The MCSymbol class represents a symbol (aka label) in the assembly file. There are two interesting kinds of symbols: assembler temporary symbols, and normal symbols. Assembler temporary symbols are used and processed by the assembler but are discarded when the object file is produced. The distinction is usually represented by adding a prefix to the label, for example “L” labels are assembler temporary labels in MachO.
MCSymbols are created by MCContext and uniqued there. This means that MCSymbols can be compared for pointer equivalence to find out if they are the same symbol. Note that pointer inequality does not guarantee the labels will end up at different addresses though. It’s perfectly legal to output something like this to the .s file:
foo:
bar:
.byte 4
In this case, both the foo and bar symbols will have the same address.
The MCSection class¶
The MCSection class represents an object-file specific section. It is
subclassed by object file specific implementations (e.g. MCSectionMachO,
MCSectionCOFF, MCSectionELF) and these are created and uniqued by
MCContext. The MCStreamer has a notion of the current section, which can be
changed with the SwitchToSection method (which corresponds to a “.section”
directive in a .s file).
The MCInst class¶
The MCInst class is a target-independent representation of an instruction.
It is a simple class (much more so than MachineInstr) that holds a
target-specific opcode and a vector of MCOperands. MCOperand, in turn, is a
simple discriminated union of three cases: 1) a simple immediate, 2) a target
register ID, 3) a symbolic expression (e.g. “Lfoo-Lbar+42”) as an MCExpr.
MCInst is the common currency used to represent machine instructions at the MC layer. It is the type used by the instruction encoder, the instruction printer, and the type generated by the assembly parser and disassembler.
Object File Format¶
The MC layer’s object writers support a variety of object formats. Because of target-specific aspects of object formats each target only supports a subset of the formats supported by the MC layer. Most targets support emitting ELF objects. Other vendor-specific objects are generally supported only on targets that are supported by that vendor (i.e. MachO is only supported on targets supported by Darwin, and XCOFF is only supported on targets that support AIX). Additionally some targets have their own object formats (i.e. DirectX, SPIR-V and WebAssembly).
The table below captures a snapshot of object file support in LLVM:
Table 104 Object File Formats¶ Format
Supported Targets
COFFAArch64, ARM, X86
DXContainerDirectX
ELFAArch64, AMDGPU, ARM, AVR, BPF, CSKY, Hexagon, Lanai, LoongArch, M86k, MSP430, MIPS, PowerPC, RISCV, SPARC, SystemZ, VE, X86
GOFFSystemZ
MachOAArch64, ARM, X86
SPIR-VSPIRV
WASMWebAssembly
XCOFFPowerPC
Target-independent code generation algorithms¶
This section documents the phases described in the high-level design of the code generator. It explains how they work and some of the rationale behind their design.
Instruction Selection¶
Instruction Selection is the process of translating LLVM code presented to the code generator into target-specific machine instructions. There are several well-known ways to do this in the literature. LLVM uses a SelectionDAG based instruction selector.
Portions of the DAG instruction selector are generated from the target
description (*.td) files. Our goal is for the entire instruction selector
to be generated from these .td files, though currently there are still
things that require custom C++ code.
GlobalISel is another instruction selection framework.
Introduction to SelectionDAGs¶
The SelectionDAG provides an abstraction for code representation in a way that is amenable to instruction selection using automatic techniques (e.g. dynamic-programming based optimal pattern matching selectors). It is also well-suited to other phases of code generation; in particular, instruction scheduling (SelectionDAG’s are very close to scheduling DAGs post-selection). Additionally, the SelectionDAG provides a host representation where a large variety of very-low-level (but target-independent) optimizations may be performed; ones which require extensive information about the instructions efficiently supported by the target.
The SelectionDAG is a Directed-Acyclic-Graph whose nodes are instances of the
SDNode class. The primary payload of the SDNode is its operation code
(Opcode) that indicates what operation the node performs and the operands to the
operation. The various operation node types are described at the top of the
include/llvm/CodeGen/ISDOpcodes.h file.
Although most operations define a single value, each node in the graph may
define multiple values. For example, a combined div/rem operation will define
both the dividend and the remainder. Many other situations require multiple
values as well. Each node also has some number of operands, which are edges to
the node defining the used value. Because nodes may define multiple values,
edges are represented by instances of the SDValue class, which is a
<SDNode, unsigned> pair, indicating the node and result value being used,
respectively. Each value produced by an SDNode has an associated MVT
(Machine Value Type) indicating what the type of the value is.
SelectionDAGs contain two different kinds of values: those that represent data
flow and those that represent control flow dependencies. Data values are simple
edges with an integer or floating point value type. Control edges are
represented as “chain” edges which are of type MVT::Other. These edges
provide an ordering between nodes that have side effects (such as loads, stores,
calls, returns, etc). All nodes that have side effects should take a token
chain as input and produce a new one as output. By convention, token chain
inputs are always operand #0, and chain results are always the last value
produced by an operation. However, after instruction selection, the
machine nodes have their chain after the instruction’s operands, and
may be followed by glue nodes.
A SelectionDAG has designated “Entry” and “Root” nodes. The Entry node is
always a marker node with an Opcode of ISD::EntryToken. The Root node is
the final side-effecting node in the token chain. For example, in a single basic
block function it would be the return node.
One important concept for SelectionDAGs is the notion of a “legal” vs. “illegal” DAG. A legal DAG for a target is one that only uses supported operations and supported types. On a 32-bit PowerPC, for example, a DAG with a value of type i1, i8, i16, or i64 would be illegal, as would a DAG that uses a SREM or UREM operation. The legalize types and legalize operations phases are responsible for turning an illegal DAG into a legal DAG.
SelectionDAG Instruction Selection Process¶
SelectionDAG-based instruction selection consists of the following steps:
Build initial DAG — This stage performs a simple translation from the input LLVM code to an illegal SelectionDAG.
Optimize SelectionDAG — This stage performs simple optimizations on the SelectionDAG to simplify it, and recognize meta instructions (like rotates and
div/rempairs) for targets that support these meta operations. This makes the resultant code more efficient and the select instructions from DAG phase (below) simpler.Legalize SelectionDAG Types — This stage transforms SelectionDAG nodes to eliminate any types that are unsupported on the target.
Optimize SelectionDAG — The SelectionDAG optimizer is run to clean up redundancies exposed by type legalization.
Legalize SelectionDAG Ops — This stage transforms SelectionDAG nodes to eliminate any operations that are unsupported on the target.
Optimize SelectionDAG — The SelectionDAG optimizer is run to eliminate inefficiencies introduced by operation legalization.
Select instructions from DAG — Finally, the target instruction selector matches the DAG operations to target instructions. This process translates the target-independent input DAG into another DAG of target instructions.
SelectionDAG Scheduling and Formation — The last phase assigns a linear order to the instructions in the target-instruction DAG and emits them into the MachineFunction being compiled. This step uses traditional prepass scheduling techniques.
After all of these steps are complete, the SelectionDAG is destroyed and the rest of the code generation passes are run.
One of the most common ways to debug these steps is using -debug-only=isel,
which prints out the DAG, along with other information like debug info,
after each of these steps. Alternatively, -debug-only=isel-dump shows only
the DAG dumps, but the results can be filtered by function names using
-filter-print-funcs=<function names>.
One great way to visualize what is going on here is to take advantage of a few LLC command line options. The following options pop up a window displaying the SelectionDAG at specific times (if you only get errors printed to the console while using this, you probably need to configure your system to add support for it).
-view-dag-combine1-dagsdisplays the DAG after being built, before the first optimization pass.-view-legalize-dagsdisplays the DAG before Legalization.-view-dag-combine2-dagsdisplays the DAG before the second optimization pass.-view-isel-dagsdisplays the DAG before the Select phase.-view-sched-dagsdisplays the DAG before Scheduling.
The -view-sunit-dags displays the Scheduler’s dependency graph. This graph
is based on the final SelectionDAG, with nodes that must be scheduled together
bundled into a single scheduling-unit node, and with immediate operands and
other nodes that aren’t relevant for scheduling omitted.
The option -filter-view-dags allows to select the name of the basic block
that you are interested to visualize and filters all the previous
view-*-dags options.
Initial SelectionDAG Construction¶
The initial SelectionDAG is naïvely peephole expanded from
the LLVM input by the SelectionDAGBuilder class. The intent of this pass
is to expose as much low-level, target-specific details to the SelectionDAG as
possible. This pass is mostly hard-coded (e.g. an LLVM add turns into an
SDNode add while a getelementptr is expanded into the obvious
arithmetic). This pass requires target-specific hooks to lower calls, returns,
varargs, etc. For these features, the TargetLowering
interface is used.
SelectionDAG LegalizeTypes Phase¶
The Legalize phase is in charge of converting a DAG to only use the types that are natively supported by the target.
There are two main ways of converting values of unsupported scalar types to values of supported types: converting small types to larger types (“promoting”), and breaking up large integer types into smaller ones (“expanding”). For example, a target might require that all f32 values are promoted to f64 and that all i1/i8/i16 values are promoted to i32. The same target might require that all i64 values be expanded into pairs of i32 values. These changes can insert sign and zero extensions as needed to make sure that the final code has the same behavior as the input.
There are two main ways of converting values of unsupported vector types to value of supported types: splitting vector types, multiple times if necessary, until a legal type is found, and extending vector types by adding elements to the end to round them out to legal types (“widening”). If a vector gets split all the way down to single-element parts with no supported vector type being found, the elements are converted to scalars (“scalarizing”).
A target implementation tells the legalizer which types are supported (and which
register class to use for them) by calling the addRegisterClass method in
its TargetLowering constructor.
SelectionDAG Legalize Phase¶
The Legalize phase is in charge of converting a DAG to only use the operations that are natively supported by the target.
Targets often have weird constraints, such as not supporting every operation on every supported datatype (e.g. X86 does not support byte conditional moves and PowerPC does not support sign-extending loads from a 16-bit memory location). Legalize takes care of this by open-coding another sequence of operations to emulate the operation (“expansion”), by promoting one type to a larger type that supports the operation (“promotion”), or by using a target-specific hook to implement the legalization (“custom”).
A target implementation tells the legalizer which operations are not supported
(and which of the above three actions to take) by calling the
setOperationAction method in its TargetLowering constructor.
If a target has legal vector types, it is expected to produce efficient machine code for common forms of the shufflevector IR instruction using those types. This may require custom legalization for SelectionDAG vector operations that are created from the shufflevector IR. The shufflevector forms that should be handled include:
Vector select — Each element of the vector is chosen from either of the corresponding elements of the 2 input vectors. This operation may also be known as a “blend” or “bitwise select” in target assembly. This type of shuffle maps directly to the
shuffle_vectorSelectionDAG node.Insert subvector — A vector is placed into a longer vector type starting at index 0. This type of shuffle maps directly to the
insert_subvectorSelectionDAG node with theindexoperand set to 0.Extract subvector — A vector is pulled from a longer vector type starting at index 0. This type of shuffle maps directly to the
extract_subvectorSelectionDAG node with theindexoperand set to 0.Splat — All elements of the vector have identical scalar elements. This operation may also be known as a “broadcast” or “duplicate” in target assembly. The shufflevector IR instruction may change the vector length, so this operation may map to multiple SelectionDAG nodes including
shuffle_vector,concat_vectors,insert_subvector, andextract_subvector.
Prior to the existence of the Legalize passes, we required that every target selector supported and handled every operator and type even if they are not natively supported. The introduction of the Legalize phases allows all of the canonicalization patterns to be shared across targets, and makes it very easy to optimize the canonicalized code because it is still in the form of a DAG.
SelectionDAG Optimization Phase: the DAG Combiner¶
The SelectionDAG optimization phase is run multiple times for code generation, immediately after the DAG is built and once after each legalization. The first run of the pass allows the initial code to be cleaned up (e.g. performing optimizations that depend on knowing that the operators have restricted type inputs). Subsequent runs of the pass clean up the messy code generated by the Legalize passes, which allows Legalize to be very simple (it can focus on making code legal instead of focusing on generating good and legal code).
One important class of optimizations performed is optimizing inserted sign and zero extension instructions. We currently use ad-hoc techniques, but could move to more rigorous techniques in the future. Here are some good papers on the subject:
“Widening integer arithmetic”
Kevin Redwine and Norman Ramsey
International Conference on Compiler Construction (CC) 2004
“Effective sign extension elimination”
Motohiro Kawahito, Hideaki Komatsu, and Toshio Nakatani
Proceedings of the ACM SIGPLAN 2002 Conference on Programming Language Design
and Implementation.
SelectionDAG Select Phase¶
The Select phase is the bulk of the target-specific code for instruction selection. This phase takes a legal SelectionDAG as input, pattern matches the instructions supported by the target to this DAG, and produces a new DAG of target code. For example, consider the following LLVM fragment:
%t1 = fadd float %W, %X
%t2 = fmul float %t1, %Y
%t3 = fadd float %t2, %Z
This LLVM code corresponds to a SelectionDAG that looks basically like this:
(fadd:f32 (fmul:f32 (fadd:f32 W, X), Y), Z)
If a target supports floating point multiply-and-add (FMA) operations, one of the adds can be merged with the multiply. On the PowerPC, for example, the output of the instruction selector might look like this DAG:
(FMADDS (FADDS W, X), Y, Z)
The FMADDS instruction is a ternary instruction that multiplies its first
two operands and adds the third (as single-precision floating-point numbers).
The FADDS instruction is a simple binary single-precision add instruction.
To perform this pattern match, the PowerPC backend includes the following
instruction definitions:
def FMADDS : AForm_1<59, 29,
(ops F4RC:$FRT, F4RC:$FRA, F4RC:$FRC, F4RC:$FRB),
"fmadds $FRT, $FRA, $FRC, $FRB",
[(set F4RC:$FRT, (fadd (fmul F4RC:$FRA, F4RC:$FRC),
F4RC:$FRB))]>;
def FADDS : AForm_2<59, 21,
(ops F4RC:$FRT, F4RC:$FRA, F4RC:$FRB),
"fadds $FRT, $FRA, $FRB",
[(set F4RC:$FRT, (fadd F4RC:$FRA, F4RC:$FRB))]>;
The highlighted portion of the instruction definitions indicates the pattern
used to match the instructions. The DAG operators (like fmul/fadd)
are defined in the include/llvm/Target/TargetSelectionDAG.td file.
“F4RC” is the register class of the input and result values.
The TableGen DAG instruction selector generator reads the instruction patterns
in the .td file and automatically builds parts of the pattern matching code
for your target. It has the following strengths:
At compiler-compile time, it analyzes your instruction patterns and tells you if your patterns make sense or not.
It can handle arbitrary constraints on operands for the pattern match. In particular, it is straight-forward to say things like “match any immediate that is a 13-bit sign-extended value”. For examples, see the
immSExt16and relatedtblgenclasses in the PowerPC backend.It knows several important identities for the patterns defined. For example, it knows that addition is commutative, so it allows the
FMADDSpattern above to match “(fadd X, (fmul Y, Z))” as well as “(fadd (fmul X, Y), Z)”, without the target author having to specially handle this case.It has a full-featured type-inferencing system. In particular, you should rarely have to explicitly tell the system what type parts of your patterns are. In the
FMADDScase above, we didn’t have to telltblgenthat all of the nodes in the pattern are of type ‘f32’. It was able to infer and propagate this knowledge from the fact thatF4RChas type ‘f32’.Targets can define their own (and rely on built-in) “pattern fragments”. Pattern fragments are chunks of reusable patterns that get inlined into your patterns during compiler-compile time. For example, the integer “
(not x)” operation is actually defined as a pattern fragment that expands as “(xor x, -1)”, since the SelectionDAG does not have a native ‘not’ operation. Targets can define their own short-hand fragments as they see fit. See the definition of ‘not’ and ‘ineg’ for examples.In addition to instructions, targets can specify arbitrary patterns that map to one or more instructions using the ‘Pat’ class. For example, the PowerPC has no way to load an arbitrary integer immediate into a register in one instruction. To tell tblgen how to do this, it defines:
// Arbitrary immediate support. Implement in terms of LIS/ORI. def : Pat<(i32 imm:$imm), (ORI (LIS (HI16 imm:$imm)), (LO16 imm:$imm))>;If none of the single-instruction patterns for loading an immediate into a register match, this will be used. This rule says “match an arbitrary i32 immediate, turning it into an
ORI(‘or a 16-bit immediate’) and anLIS(‘load 16-bit immediate, where the immediate is shifted to the left 16 bits’) instruction”. To make this work, theLO16/HI16node transformations are used to manipulate the input immediate (in this case, take the high or low 16-bits of the immediate).When using the ‘Pat’ class to map a pattern to an instruction that has one or more complex operands (like e.g. X86 addressing mode), the pattern may either specify the operand as a whole using a
ComplexPattern, or else it may specify the components of the complex operand separately. The latter is done e.g. for pre-increment instructions by the PowerPC back end:def STWU : DForm_1<37, (outs ptr_rc:$ea_res), (ins GPRC:$rS, memri:$dst), "stwu $rS, $dst", LdStStoreUpd, []>, RegConstraint<"$dst.reg = $ea_res">, NoEncode<"$ea_res">; def : Pat<(pre_store GPRC:$rS, ptr_rc:$ptrreg, iaddroff:$ptroff), (STWU GPRC:$rS, iaddroff:$ptroff, ptr_rc:$ptrreg)>;Here, the pair of
ptroffandptrregoperands is matched onto the complex operanddstof classmemriin theSTWUinstruction.While the system does automate a lot, it still allows you to write custom C++ code to match special cases if there is something that is hard to express.
While it has many strengths, the system currently has some limitations, primarily because it is a work in progress and is not yet finished:
Overall, there is no way to define or match SelectionDAG nodes that define multiple values (e.g.
SMUL_LOHI,LOAD,CALL, etc). This is the biggest reason that you currently still have to write custom C++ code for your instruction selector.There is no great way to support matching complex addressing modes yet. In the future, we will extend pattern fragments to allow them to define multiple values (e.g. the four operands of the X86 addressing mode, which are currently matched with custom C++ code). In addition, we’ll extend fragments so that a fragment can match multiple different patterns.
We don’t automatically infer flags like
isStore/isLoadyet.We don’t automatically generate the set of supported registers and operations for the Legalizer yet.
We don’t have a way of tying in custom legalized nodes yet.
Despite these limitations, the instruction selector generator is still quite useful for most of the binary and logical operations in typical instruction sets. If you run into any problems or can’t figure out how to do something, please let Chris know!
SelectionDAG Scheduling and Formation Phase¶
The scheduling phase takes the DAG of target instructions from the selection phase and assigns an order. The scheduler can pick an order depending on various constraints of the machines (i.e. order for minimal register pressure or try to cover instruction latencies). Once an order is established, the DAG is converted to a list of MachineInstrs and the SelectionDAG is destroyed.
Note that this phase is logically separate from the instruction selection phase, but is tied to it closely in the code because it operates on SelectionDAGs.
Future directions for the SelectionDAG¶
Optional function-at-a-time selection.
Auto-generate entire selector from
.tdfile.
SSA-based Machine Code Optimizations¶
To Be Written
Live Intervals¶
Live Intervals are the ranges (intervals) where a variable is live. They are used by some register allocator passes to determine if two or more virtual registers which require the same physical register are live at the same point in the program (i.e., they conflict). When this situation occurs, one virtual register must be spilled.
Live Variable Analysis¶
The first step in determining the live intervals of variables is to calculate the set of registers that are immediately dead after the instruction (i.e., the instruction calculates the value, but it is never used) and the set of registers that are used by the instruction, but are never used after the instruction (i.e., they are killed). Live variable information is computed for each virtual register and register allocatable physical register in the function. This is done in a very efficient manner because it uses SSA to sparsely compute lifetime information for virtual registers (which are in SSA form) and only has to track physical registers within a block. Before register allocation, LLVM can assume that physical registers are only live within a single basic block. This allows it to do a single, local analysis to resolve physical register lifetimes within each basic block. If a physical register is not register allocatable (e.g., a stack pointer or condition codes), it is not tracked.
Physical registers may be live in to or out of a function. Live in values are
typically arguments in registers. Live out values are typically return values in
registers. Live in values are marked as such, and are given a dummy “defining”
instruction during live intervals analysis. If the last basic block of a
function is a return, then it’s marked as using all live out values in the
function.
PHI nodes need to be handled specially, because the calculation of the live
variable information from a depth first traversal of the CFG of the function
won’t guarantee that a virtual register used by the PHI node is defined
before it’s used. When a PHI node is encountered, only the definition is
handled, because the uses will be handled in other basic blocks.
For each PHI node of the current basic block, we simulate an assignment at
the end of the current basic block and traverse the successor basic blocks. If a
successor basic block has a PHI node and one of the PHI node’s operands
is coming from the current basic block, then the variable is marked as alive
within the current basic block and all of its predecessor basic blocks, until
the basic block with the defining instruction is encountered.
Live Intervals Analysis¶
We now have the information available to perform the live intervals analysis and
build the live intervals themselves. We start off by numbering the basic blocks
and machine instructions. We then handle the “live-in” values. These are in
physical registers, so the physical register is assumed to be killed by the end
of the basic block. Live intervals for virtual registers are computed for some
ordering of the machine instructions [1, N]. A live interval is an interval
[i, j), where 1 >= i >= j > N, for which a variable is live.
Note
More to come…
Register Allocation¶
The Register Allocation problem consists in mapping a program Pv, that can use an unbounded number of virtual registers, to a program Pp that contains a finite (possibly small) number of physical registers. Each target architecture has a different number of physical registers. If the number of physical registers is not enough to accommodate all the virtual registers, some of them will have to be mapped into memory. These virtuals are called spilled virtuals.
How registers are represented in LLVM¶
In LLVM, physical registers are denoted by integer numbers that normally range
from 1 to 1023. To see how this numbering is defined for a particular
architecture, you can read the GenRegisterNames.inc file for that
architecture. For instance, by inspecting
lib/Target/X86/X86GenRegisterInfo.inc we see that the 32-bit register
EAX is denoted by 43, and the MMX register MM0 is mapped to 65.
Some architectures contain registers that share the same physical location. A
notable example is the X86 platform. For instance, in the X86 architecture, the
registers EAX, AX and AL share the first eight bits. These physical
registers are marked as aliased in LLVM. Given a particular architecture, you
can check which registers are aliased by inspecting its RegisterInfo.td
file. Moreover, the class MCRegAliasIterator enumerates all the physical
registers aliased to a register.
Physical registers, in LLVM, are grouped in Register Classes. Elements in the
same register class are functionally equivalent, and can be interchangeably
used. Each virtual register can only be mapped to physical registers of a
particular class. For instance, in the X86 architecture, some virtuals can only
be allocated to 8 bit registers. A register class is described by
TargetRegisterClass objects. To discover if a virtual register is
compatible with a given physical, this code can be used:
bool RegMapping_Fer::compatible_class(MachineFunction &mf,
unsigned v_reg,
unsigned p_reg) {
assert(TargetRegisterInfo::isPhysicalRegister(p_reg) &&
"Target register must be physical");
const TargetRegisterClass *trc = mf.getRegInfo().getRegClass(v_reg);
return trc->contains(p_reg);
}
Sometimes, mostly for debugging purposes, it is useful to change the number of
physical registers available in the target architecture. This must be done
statically, inside the TargetRegisterInfo.td file. Just grep for
RegisterClass, the last parameter of which is a list of registers. Just
commenting some out is one simple way to avoid them being used. A more polite
way is to explicitly exclude some registers from the allocation order. See the
definition of the GR8 register class in
lib/Target/X86/X86RegisterInfo.td for an example of this.
Virtual registers are also denoted by integer numbers. Contrary to physical
registers, different virtual registers never share the same number. Whereas
physical registers are statically defined in a TargetRegisterInfo.td file
and cannot be created by the application developer, that is not the case with
virtual registers. In order to create new virtual registers, use the method
MachineRegisterInfo::createVirtualRegister(). This method will return a new
virtual register. Use an IndexedMap<Foo, VirtReg2IndexFunctor> to hold
information per virtual register. If you need to enumerate all virtual
registers, use the function TargetRegisterInfo::index2VirtReg() to find the
virtual register numbers:
for (unsigned i = 0, e = MRI->getNumVirtRegs(); i != e; ++i) {
unsigned VirtReg = TargetRegisterInfo::index2VirtReg(i);
stuff(VirtReg);
}
Before register allocation, the operands of an instruction are mostly virtual
registers, although physical registers may also be used. In order to check if a
given machine operand is a register, use the boolean function
MachineOperand::isRegister(). To obtain the integer code of a register, use
MachineOperand::getReg(). An instruction may define or use a register. For
instance, ADD reg:1026 := reg:1025 reg:1024 defines the registers 1024, and
uses registers 1025 and 1026. Given a register operand, the method
MachineOperand::isUse() informs if that register is being used by the
instruction. The method MachineOperand::isDef() informs if that registers is
being defined.
We will call physical registers present in the LLVM bitcode before register
allocation pre-colored registers. Pre-colored registers are used in many
different situations, for instance, to pass parameters of functions calls, and
to store results of particular instructions. There are two types of pre-colored
registers: the ones implicitly defined, and those explicitly
defined. Explicitly defined registers are normal operands, and can be accessed
with MachineInstr::getOperand(int)::getReg(). In order to check which
registers are implicitly defined by an instruction, use the
TargetInstrInfo::get(opcode)::ImplicitDefs, where opcode is the opcode
of the target instruction. One important difference between explicit and
implicit physical registers is that the latter are defined statically for each
instruction, whereas the former may vary depending on the program being
compiled. For example, an instruction that represents a function call will
always implicitly define or use the same set of physical registers. To read the
registers implicitly used by an instruction, use
TargetInstrInfo::get(opcode)::ImplicitUses. Pre-colored registers impose
constraints on any register allocation algorithm. The register allocator must
make sure that none of them are overwritten by the values of virtual registers
while still alive.
Mapping virtual registers to physical registers¶
There are two ways to map virtual registers to physical registers (or to memory
slots). The first way, that we will call direct mapping, is based on the use
of methods of the classes TargetRegisterInfo, and MachineOperand. The
second way, that we will call indirect mapping, relies on the VirtRegMap
class in order to insert loads and stores sending and getting values to and from
memory.
The direct mapping provides more flexibility to the developer of the register
allocator; however, it is more error prone, and demands more implementation
work. Basically, the programmer will have to specify where load and store
instructions should be inserted in the target function being compiled in order
to get and store values in memory. To assign a physical register to a virtual
register present in a given operand, use MachineOperand::setReg(p_reg). To
insert a store instruction, use TargetInstrInfo::storeRegToStackSlot(...),
and to insert a load instruction, use TargetInstrInfo::loadRegFromStackSlot.
The indirect mapping shields the application developer from the complexities of
inserting load and store instructions. In order to map a virtual register to a
physical one, use VirtRegMap::assignVirt2Phys(vreg, preg). In order to map
a certain virtual register to memory, use
VirtRegMap::assignVirt2StackSlot(vreg). This method will return the stack
slot where vreg’s value will be located. If it is necessary to map another
virtual register to the same stack slot, use
VirtRegMap::assignVirt2StackSlot(vreg, stack_location). One important point
to consider when using the indirect mapping, is that even if a virtual register
is mapped to memory, it still needs to be mapped to a physical register. This
physical register is the location where the virtual register is supposed to be
found before being stored or after being reloaded.
If the indirect strategy is used, after all the virtual registers have been
mapped to physical registers or stack slots, it is necessary to use a spiller
object to place load and store instructions in the code. Every virtual that has
been mapped to a stack slot will be stored to memory after being defined and will
be loaded before being used. The implementation of the spiller tries to recycle
load/store instructions, avoiding unnecessary instructions. For an example of
how to invoke the spiller, see RegAllocLinearScan::runOnMachineFunction in
lib/CodeGen/RegAllocLinearScan.cpp.
Handling two address instructions¶
With very rare exceptions (e.g., function calls), the LLVM machine code
instructions are three address instructions. That is, each instruction is
expected to define at most one register, and to use at most two registers.
However, some architectures use two address instructions. In this case, the
defined register is also one of the used registers. For instance, an instruction
such as ADD %EAX, %EBX, in X86 is actually equivalent to %EAX = %EAX +
%EBX.
In order to produce correct code, LLVM must convert three address instructions
that represent two address instructions into true two address instructions. LLVM
provides the pass TwoAddressInstructionPass for this specific purpose. It
must be run before register allocation takes place. After its execution, the
resulting code may no longer be in SSA form. This happens, for instance, in
situations where an instruction such as %a = ADD %b %c is converted to two
instructions such as:
%a = MOVE %b
%a = ADD %a %c
Notice that, internally, the second instruction is represented as ADD
%a[def/use] %c. I.e., the register operand %a is both used and defined by
the instruction.
The SSA deconstruction phase¶
An important transformation that happens during register allocation is called the SSA Deconstruction Phase. The SSA form simplifies many analyses that are performed on the control flow graph of programs. However, traditional instruction sets do not implement PHI instructions. Thus, in order to generate executable code, compilers must replace PHI instructions with other instructions that preserve their semantics.
There are many ways in which PHI instructions can safely be removed from the
target code. The most traditional PHI deconstruction algorithm replaces PHI
instructions with copy instructions. That is the strategy adopted by LLVM. The
SSA deconstruction algorithm is implemented in
lib/CodeGen/PHIElimination.cpp. In order to invoke this pass, the identifier
PHIEliminationID must be marked as required in the code of the register
allocator.
Instruction folding¶
Instruction folding is an optimization performed during register allocation that removes unnecessary copy instructions. For instance, a sequence of instructions such as:
%EBX = LOAD %mem_address
%EAX = COPY %EBX
can be safely substituted by the single instruction:
%EAX = LOAD %mem_address
Instructions can be folded with the
TargetRegisterInfo::foldMemoryOperand(...) method. Care must be taken when
folding instructions; a folded instruction can be quite different from the
original instruction. See LiveIntervals::addIntervalsForSpills in
lib/CodeGen/LiveIntervalAnalysis.cpp for an example of its use.
Built in register allocators¶
The LLVM infrastructure provides the application developer with three different register allocators:
Fast — This register allocator is the default for debug builds. It allocates registers on a basic block level, attempting to keep values in registers and reusing registers as appropriate.
Basic — This is an incremental approach to register allocation. Live ranges are assigned to registers one at a time in an order that is driven by heuristics. Since code can be rewritten on-the-fly during allocation, this framework allows interesting allocators to be developed as extensions. It is not itself a production register allocator but is a potentially useful stand-alone mode for triaging bugs and as a performance baseline.
Greedy — The default allocator. This is a highly tuned implementation of the Basic allocator that incorporates global live range splitting. This allocator works hard to minimize the cost of spill code.
PBQP — A Partitioned Boolean Quadratic Programming (PBQP) based register allocator. This allocator works by constructing a PBQP problem representing the register allocation problem under consideration, solving this using a PBQP solver, and mapping the solution back to a register assignment.
The type of register allocator used in llc can be chosen with the command
line option -regalloc=...:
$ llc -regalloc=linearscan file.bc -o ln.s
$ llc -regalloc=fast file.bc -o fa.s
$ llc -regalloc=pbqp file.bc -o pbqp.s
Prolog/Epilog Code Insertion¶
Note
To Be Written
Compact Unwind¶
Throwing an exception requires unwinding out of a function. The information on how to unwind a given function is traditionally expressed in DWARF unwind (a.k.a. frame) info. But that format was originally developed for debuggers to backtrace, and each Frame Description Entry (FDE) requires ~20-30 bytes per function. There is also the cost of mapping from an address in a function to the corresponding FDE at runtime. An alternative unwind encoding is called compact unwind and requires just 4-bytes per function.
The compact unwind encoding is a 32-bit value, which is encoded in an
architecture-specific way. It specifies which registers to restore and from
where, and how to unwind out of the function. When the linker creates a final
linked image, it will create a __TEXT,__unwind_info section. This section is
a small and fast way for the runtime to access unwind info for any given
function. If we emit compact unwind info for the function, that compact unwind
info will be encoded in the __TEXT,__unwind_info section. If we emit DWARF
unwind info, the __TEXT,__unwind_info section will contain the offset of the
FDE in the __TEXT,__eh_frame section in the final linked image.
For X86, there are three modes for the compact unwind encoding:
- Function with a Frame Pointer (``EBP`` or ``RBP``)
EBP/RBP-based frame, whereEBP/RBPis pushed onto the stack immediately after the return address, thenESP/RSPis moved toEBP/RBP. Thus to unwind,ESP/RSPis restored with the currentEBP/RBPvalue, thenEBP/RBPis restored by popping the stack, and the return is done by popping the stack once more into the PC. All non-volatile registers that need to be restored must have been saved in a small range on the stack that startsEBP-4toEBP-1020(RBP-8toRBP-1020). The offset (divided by 4 in 32-bit mode and 8 in 64-bit mode) is encoded in bits 16-23 (mask:0x00FF0000). The registers saved are encoded in bits 0-14 (mask:0x00007FFF) as five 3-bit entries from the following table:Compact Number
i386 Register
x86-64 Register
1
EBXRBX2
ECXR123
EDXR134
EDIR145
ESIR156
EBPRBP- Frameless with a Small Constant Stack Size (``EBP`` or ``RBP`` is not used as a frame pointer)
To return, a constant (encoded in the compact unwind encoding) is added to the
ESP/RSP. Then the return is done by popping the stack into the PC. All non-volatile registers that need to be restored must have been saved on the stack immediately after the return address. The stack size (divided by 4 in 32-bit mode and 8 in 64-bit mode) is encoded in bits 16-23 (mask:0x00FF0000). There is a maximum stack size of 1024 bytes in 32-bit mode and 2048 in 64-bit mode. The number of registers saved is encoded in bits 9-12 (mask:0x00001C00). Bits 0-9 (mask:0x000003FF) contain which registers were saved and their order. (See theencodeCompactUnwindRegistersWithoutFrame()function inlib/Target/X86FrameLowering.cppfor the encoding algorithm.)- Frameless with a Large Constant Stack Size (``EBP`` or ``RBP`` is not used as a frame pointer)
This case is like the “Frameless with a Small Constant Stack Size” case, but the stack size is too large to encode in the compact unwind encoding. Instead it requires that the function contains “
subl $nnnnnn, %esp” in its prolog. The compact encoding contains the offset to the$nnnnnnvalue in the function in bits 9-12 (mask:0x00001C00).
Late Machine Code Optimizations¶
Note
To Be Written
Code Emission¶
The code emission step of code generation is responsible for lowering from the code generator abstractions (like MachineFunction, MachineInstr, etc) down to the abstractions used by the MC layer (MCInst, MCStreamer, etc). This is done with a combination of several different classes: the (misnamed) target-independent AsmPrinter class, target-specific subclasses of AsmPrinter (such as SparcAsmPrinter), and the TargetLoweringObjectFile class.
Since the MC layer works at the level of abstraction of object files, it doesn’t have a notion of functions, global variables etc. Instead, it thinks about labels, directives, and instructions. A key class used at this time is the MCStreamer class. This is an abstract API that is implemented in different ways (e.g. to output a .s file, output an ELF .o file, etc) that is effectively an “assembler API”. MCStreamer has one method per directive, such as EmitLabel, EmitSymbolAttribute, switchSection, etc, which directly correspond to assembly level directives.
If you are interested in implementing a code generator for a target, there are three important things that you have to implement for your target:
First, you need a subclass of AsmPrinter for your target. This class implements the general lowering process converting MachineFunction’s into MC label constructs. The AsmPrinter base class provides a number of useful methods and routines, and also allows you to override the lowering process in some important ways. You should get much of the lowering for free if you are implementing an ELF, COFF, or MachO target, because the TargetLoweringObjectFile class implements much of the common logic.
Second, you need to implement an instruction printer for your target. The instruction printer takes an MCInst and renders it to a raw_ostream as text. Most of this is automatically generated from the .td file (when you specify something like “
add $dst, $src1, $src2” in the instructions), but you need to implement routines to print operands.Third, you need to implement code that lowers a MachineInstr to an MCInst, usually implemented in “<target>MCInstLower.cpp”. This lowering process is often target specific, and is responsible for turning jump table entries, constant pool indices, global variable addresses, etc into MCLabels as appropriate. This translation layer is also responsible for expanding pseudo ops used by the code generator into the actual machine instructions they correspond to. The MCInsts that are generated by this are fed into the instruction printer or the encoder.
Finally, at your choosing, you can also implement a subclass of MCCodeEmitter which lowers MCInst’s into machine code bytes and relocations. This is important if you want to support direct .o file emission, or would like to implement an assembler for your target.
Emitting function stack size information¶
A section containing metadata on function stack sizes will be emitted when
TargetLoweringObjectFile::StackSizesSection is not null, and
TargetOptions::EmitStackSizeSection is set (-stack-size-section). The
section will contain an array of pairs of function symbol values (pointer size)
and stack sizes (unsigned LEB128). The stack size values only include the space
allocated in the function prologue. Functions with dynamic stack allocations are
not included.
VLIW Packetizer¶
In a Very Long Instruction Word (VLIW) architecture, the compiler is responsible for mapping instructions to functional-units available on the architecture. To that end, the compiler creates groups of instructions called packets or bundles. The VLIW packetizer in LLVM is a target-independent mechanism to enable the packetization of machine instructions.
Mapping from instructions to functional units¶
Instructions in a VLIW target can typically be mapped to multiple functional units. During the process of packetizing, the compiler must be able to reason about whether an instruction can be added to a packet. This decision can be complex since the compiler has to examine all possible mappings of instructions to functional units. Therefore to alleviate compilation-time complexity, the VLIW packetizer parses the instruction classes of a target and generates tables at compiler build time. These tables can then be queried by the provided machine-independent API to determine if an instruction can be accommodated in a packet.
How the packetization tables are generated and used¶
The packetizer reads instruction classes from a target’s itineraries and creates a deterministic finite automaton (DFA) to represent the state of a packet. A DFA consists of three major elements: inputs, states, and transitions. The set of inputs for the generated DFA represents the instruction being added to a packet. The states represent the possible consumption of functional units by instructions in a packet. In the DFA, transitions from one state to another occur on the addition of an instruction to an existing packet. If there is a legal mapping of functional units to instructions, then the DFA contains a corresponding transition. The absence of a transition indicates that a legal mapping does not exist and that the instruction cannot be added to the packet.
To generate tables for a VLIW target, add TargetGenDFAPacketizer.inc as a
target to the Makefile in the target directory. The exported API provides three
functions: DFAPacketizer::clearResources(),
DFAPacketizer::reserveResources(MachineInstr *MI), and
DFAPacketizer::canReserveResources(MachineInstr *MI). These functions allow
a target packetizer to add an instruction to an existing packet and to check
whether an instruction can be added to a packet. See
llvm/CodeGen/DFAPacketizer.h for more information.
Implementing a Native Assembler¶
Though you’re probably reading this because you want to write or maintain a compiler backend, LLVM also fully supports building a native assembler. We’ve tried hard to automate the generation of the assembler from the .td files (in particular the instruction syntax and encodings), which means that a large part of the manual and repetitive data entry can be factored and shared with the compiler.
Instruction Parsing¶
Note
To Be Written
Instruction Alias Processing¶
Once the instruction is parsed, it enters the MatchInstructionImpl function. The MatchInstructionImpl function performs alias processing and then does actual matching.
Alias processing is the phase that canonicalizes different lexical forms of the same instructions down to one representation. There are several different kinds of alias that are possible to implement and they are listed below in the order that they are processed (which is in order from simplest/weakest to most complex/powerful). Generally you want to use the first alias mechanism that meets the needs of your instruction, because it will allow a more concise description.
Mnemonic Aliases¶
The first phase of alias processing is simple instruction mnemonic remapping for classes of instructions which are allowed with two different mnemonics. This phase is a simple and unconditionally remapping from one input mnemonic to one output mnemonic. It isn’t possible for this form of alias to look at the operands at all, so the remapping must apply for all forms of a given mnemonic. Mnemonic aliases are defined simply, for example X86 has:
def : MnemonicAlias<"cbw", "cbtw">;
def : MnemonicAlias<"smovq", "movsq">;
def : MnemonicAlias<"fldcww", "fldcw">;
def : MnemonicAlias<"fucompi", "fucomip">;
def : MnemonicAlias<"ud2a", "ud2">;
… and many others. With a MnemonicAlias definition, the mnemonic is remapped simply and directly. Though MnemonicAlias’s can’t look at any aspect of the instruction (such as the operands) they can depend on global modes (the same ones supported by the matcher), through a Requires clause:
def : MnemonicAlias<"pushf", "pushfq">, Requires<[In64BitMode]>;
def : MnemonicAlias<"pushf", "pushfl">, Requires<[In32BitMode]>;
In this example, the mnemonic gets mapped into a different one depending on the current instruction set.
Instruction Aliases¶
The most general phase of alias processing occurs while matching is happening: it provides new forms for the matcher to match along with a specific instruction to generate. An instruction alias has two parts: the string to match and the instruction to generate. For example:
def : InstAlias<"movsx $src, $dst", (MOVSX16rr8W GR16:$dst, GR8 :$src)>;
def : InstAlias<"movsx $src, $dst", (MOVSX16rm8W GR16:$dst, i8mem:$src)>;
def : InstAlias<"movsx $src, $dst", (MOVSX32rr8 GR32:$dst, GR8 :$src)>;
def : InstAlias<"movsx $src, $dst", (MOVSX32rr16 GR32:$dst, GR16 :$src)>;
def : InstAlias<"movsx $src, $dst", (MOVSX64rr8 GR64:$dst, GR8 :$src)>;
def : InstAlias<"movsx $src, $dst", (MOVSX64rr16 GR64:$dst, GR16 :$src)>;
def : InstAlias<"movsx $src, $dst", (MOVSX64rr32 GR64:$dst, GR32 :$src)>;
This shows a powerful example of the instruction aliases, matching the same mnemonic in multiple different ways depending on what operands are present in the assembly. The result of instruction aliases can include operands in a different order than the destination instruction, and can use an input multiple times, for example:
def : InstAlias<"clrb $reg", (XOR8rr GR8 :$reg, GR8 :$reg)>;
def : InstAlias<"clrw $reg", (XOR16rr GR16:$reg, GR16:$reg)>;
def : InstAlias<"clrl $reg", (XOR32rr GR32:$reg, GR32:$reg)>;
def : InstAlias<"clrq $reg", (XOR64rr GR64:$reg, GR64:$reg)>;
This example also shows that tied operands are only listed once. In the X86 backend, XOR8rr has two input GR8’s and one output GR8 (where an input is tied to the output). InstAliases take a flattened operand list without duplicates for tied operands. The result of an instruction alias can also use immediates and fixed physical registers which are added as simple immediate operands in the result, for example:
// Fixed Immediate operand.
def : InstAlias<"aad", (AAD8i8 10)>;
// Fixed register operand.
def : InstAlias<"fcomi", (COM_FIr ST1)>;
// Simple alias.
def : InstAlias<"fcomi $reg", (COM_FIr RST:$reg)>;
Instruction aliases can also have a Requires clause to make them subtarget specific.
If the back-end supports it, the instruction printer can automatically emit the alias rather than what’s being aliased. It typically leads to better, more readable code. If it’s better to print out what’s being aliased, then pass a ‘0’ as the third parameter to the InstAlias definition.
Instruction Matching¶
Note
To Be Written
Target-specific Implementation Notes¶
This section of the document explains features or design decisions that are specific to the code generator for a particular target.
Tail call optimization¶
Tail call optimization, callee reusing the stack of the caller, is currently supported on x86/x86-64, PowerPC, AArch64, and WebAssembly. It is performed on x86/x86-64, PowerPC, and AArch64 if:
Caller and callee have the calling convention
fastcc,cc 10(GHC calling convention),cc 11(HiPE calling convention),tailcc, orswifttailcc.The call is a tail call - in tail position (ret immediately follows call and ret uses value of call or is void).
Option
-tailcalloptis enabled or the calling convention istailcc.Platform-specific constraints are met.
x86/x86-64 constraints:
No variable argument lists are used.
On x86-64 when generating GOT/PIC code only module-local calls (visibility = hidden or protected) are supported.
PowerPC constraints:
No variable argument lists are used.
No byval parameters are used.
On ppc32/64 GOT/PIC only module-local calls (visibility = hidden or protected) are supported.
WebAssembly constraints:
No variable argument lists are used
The ‘tail-call’ target attribute is enabled.
The caller and callee’s return types must match. The caller cannot be void unless the callee is, too.
AArch64 constraints:
No variable argument lists are used.
Example:
Call as llc -tailcallopt test.ll.
declare fastcc i32 @tailcallee(i32 inreg %a1, i32 inreg %a2, i32 %a3, i32 %a4)
define fastcc i32 @tailcaller(i32 %in1, i32 %in2) {
%l1 = add i32 %in1, %in2
%tmp = tail call fastcc i32 @tailcallee(i32 inreg %in1, i32 inreg %in2, i32 %in1, i32 %l1)
ret i32 %tmp
}
Implications of -tailcallopt:
To support tail call optimization in situations where the callee has more
arguments than the caller a ‘callee pops arguments’ convention is used. This
currently causes each fastcc call that is not tail call optimized (because
one or more of above constraints are not met) to be followed by a readjustment
of the stack. So performance might be worse in such cases.
Sibling call optimization¶
Sibling call optimization is a restricted form of tail call optimization.
Unlike tail call optimization described in the previous section, it can be
performed automatically on any tail calls when -tailcallopt option is not
specified.
Sibling call optimization is currently performed on x86/x86-64 when the following constraints are met:
Caller and callee have the same calling convention. It can be either
corfastcc.The call is a tail call - in tail position (ret immediately follows call and ret uses value of call or is void).
Caller and callee have matching return type or the callee result is not used.
If any of the callee arguments are being passed in stack, they must be available in caller’s own incoming argument stack and the frame offsets must be the same.
Example:
declare i32 @bar(i32, i32)
define i32 @foo(i32 %a, i32 %b, i32 %c) {
entry:
%0 = tail call i32 @bar(i32 %a, i32 %b)
ret i32 %0
}
The X86 backend¶
The X86 code generator lives in the lib/Target/X86 directory. This code
generator is capable of targeting a variety of x86-32 and x86-64 processors, and
includes support for ISA extensions such as MMX and SSE.
X86 Target Triples supported¶
The following are the known target triples that are supported by the X86 backend. This is not an exhaustive list, and it would be useful to add those that people test.
i686-pc-linux-gnu — Linux
i386-unknown-freebsd5.3 — FreeBSD 5.3
i686-pc-cygwin — Cygwin on Win32
i686-pc-mingw32 — MingW on Win32
i386-pc-mingw32msvc — MingW crosscompiler on Linux
i686-apple-darwin* — Apple Darwin on X86
x86_64-unknown-linux-gnu — Linux
X86 Calling Conventions supported¶
The following target-specific calling conventions are known to backend:
x86_StdCall — stdcall calling convention seen on Microsoft Windows platform (CC ID = 64).
x86_FastCall — fastcall calling convention seen on Microsoft Windows platform (CC ID = 65).
x86_ThisCall — Similar to X86_StdCall. Passes first argument in ECX, others via stack. Callee is responsible for stack cleaning. This convention is used by MSVC by default for methods in its ABI (CC ID = 70).
Representing X86 addressing modes in MachineInstrs¶
The x86 has a very flexible way of accessing memory. It is capable of forming memory addresses of the following expression directly in integer instructions (which use ModR/M addressing):
SegmentReg: Base + [1,2,4,8] * IndexReg + Disp32
In order to represent this, LLVM tracks no less than 5 operands for each memory
operand of this form. This means that the “load” form of ‘mov’ has the
following MachineOperands in this order:
Index: 0 | 1 2 3 4 5
Meaning: DestReg, | BaseReg, Scale, IndexReg, Displacement Segment
OperandTy: VirtReg, | VirtReg, UnsImm, VirtReg, SignExtImm PhysReg
Stores, and all other instructions, treat the four memory operands in the same way and in the same order. If the segment register is unspecified (regno = 0), then no segment override is generated. “Lea” operations do not have a segment register specified, so they only have 4 operands for their memory reference.
X86 address spaces supported¶
x86 has a feature which provides the ability to perform loads and stores to different address spaces via the x86 segment registers. A segment override prefix byte on an instruction causes the instruction’s memory access to go to the specified segment. LLVM address space 0 is the default address space, which includes the stack, and any unqualified memory accesses in a program. Address spaces 1-255 are currently reserved for user-defined code. The GS-segment is represented by address space 256, the FS-segment is represented by address space 257, and the SS-segment is represented by address space 258. Other x86 segments have yet to be allocated address space numbers.
While these address spaces may seem similar to TLS via the thread_local
keyword, and often use the same underlying hardware, there are some fundamental
differences.
The thread_local keyword applies to global variables and specifies that they
are to be allocated in thread-local memory. There are no type qualifiers
involved, and these variables can be pointed to with normal pointers and
accessed with normal loads and stores. The thread_local keyword is
target-independent at the LLVM IR level (though LLVM doesn’t yet have
implementations of it for some configurations)
Special address spaces, in contrast, apply to static types. Every load and store has a particular address space in its address operand type, and this is what determines which address space is accessed. LLVM ignores these special address space qualifiers on global variables, and does not provide a way to directly allocate storage in them. At the LLVM IR level, the behavior of these special address spaces depends in part on the underlying OS or runtime environment, and they are specific to x86 (and LLVM doesn’t yet handle them correctly in some cases).
Some operating systems and runtime environments use (or may in the future use) the FS/GS-segment registers for various low-level purposes, so care should be taken when considering them.
Instruction naming¶
An instruction name consists of the base name, a default operand size, and a character per operand with an optional special size. For example:
ADD8rr -> add, 8-bit register, 8-bit register
IMUL16rmi -> imul, 16-bit register, 16-bit memory, 16-bit immediate
IMUL16rmi8 -> imul, 16-bit register, 16-bit memory, 8-bit immediate
MOVSX32rm16 -> movsx, 32-bit register, 16-bit memory
The PowerPC backend¶
The PowerPC code generator lives in the lib/Target/PowerPC directory. The code generation is retargetable to several variations or subtargets of the PowerPC ISA; including ppc32, ppc64 and altivec.
LLVM PowerPC ABI¶
LLVM follows the AIX PowerPC ABI, with two deviations. LLVM uses a PC relative (PIC) or static addressing for accessing global values, so no TOC (r2) is used. Second, r31 is used as a frame pointer to allow dynamic growth of a stack frame. LLVM takes advantage of having no TOC to provide space to save the frame pointer in the PowerPC linkage area of the caller frame. Other details of PowerPC ABI can be found at PowerPC ABI. Note: This link describes the 32 bit ABI. The 64 bit ABI is similar except space for GPRs are 8 bytes wide (not 4) and r13 is reserved for system use.
Frame Layout¶
The size of a PowerPC frame is usually fixed for the duration of a function’s invocation. Since the frame is fixed size, all references into the frame can be accessed via fixed offsets from the stack pointer. The exception to this is when dynamic alloca or variable sized arrays are present, then a base pointer (r31) is used as a proxy for the stack pointer and stack pointer is free to grow or shrink. A base pointer is also used if llvm-gcc is not passed the -fomit-frame-pointer flag. The stack pointer is always aligned to 16 bytes, so that space allocated for altivec vectors will be properly aligned.
An invocation frame is laid out as follows (low memory at top):
Linkage
Parameter area
Dynamic area
Locals area
Saved registers area
Previous Frame
The linkage area is used by a callee to save special registers prior to allocating its own frame. Only three entries are relevant to LLVM. The first entry is the previous stack pointer (sp), aka link. This allows probing tools like gdb or exception handlers to quickly scan the frames in the stack. A function epilog can also use the link to pop the frame from the stack. The third entry in the linkage area is used to save the return address from the lr register. Finally, as mentioned above, the last entry is used to save the previous frame pointer (r31.) The entries in the linkage area are the size of a GPR, thus the linkage area is 24 bytes long in 32 bit mode and 48 bytes in 64 bit mode.
32 bit linkage area:
0
Saved SP (r1)
4
Saved CR
8
Saved LR
12
Reserved
16
Reserved
20
Saved FP (r31)
64 bit linkage area:
0
Saved SP (r1)
8
Saved CR
16
Saved LR
24
Reserved
32
Reserved
40
Saved FP (r31)
The parameter area is used to store arguments being passed to a callee function. Following the PowerPC ABI, the first few arguments are actually passed in registers, with the space in the parameter area unused. However, if there are not enough registers or the callee is a thunk or vararg function, these register arguments can be spilled into the parameter area. Thus, the parameter area must be large enough to store all the parameters for the largest call sequence made by the caller. The size must also be minimally large enough to spill registers r3-r10. This allows callees blind to the call signature, such as thunks and vararg functions, enough space to cache the argument registers. Therefore, the parameter area is minimally 32 bytes (64 bytes in 64 bit mode.) Also note that since the parameter area is a fixed offset from the top of the frame, that a callee can access its split arguments using fixed offsets from the stack pointer (or base pointer.)
Combining the information about the linkage, parameter areas and alignment. A stack frame is minimally 64 bytes in 32 bit mode and 128 bytes in 64 bit mode.
The dynamic area starts out as size zero. If a function uses dynamic alloca then space is added to the stack, the linkage and parameter areas are shifted to top of stack, and the new space is available immediately below the linkage and parameter areas. The cost of shifting the linkage and parameter areas is minor since only the link value needs to be copied. The link value can be easily fetched by adding the original frame size to the base pointer. Note that allocations in the dynamic space need to observe 16 byte alignment.
The locals area is where the llvm compiler reserves space for local variables.
The saved registers area is where the llvm compiler spills callee saved registers on entry to the callee.
Prolog/Epilog¶
The llvm prolog and epilog are the same as described in the PowerPC ABI, with the following exceptions. Callee saved registers are spilled after the frame is created. This allows the llvm epilog/prolog support to be common with other targets. The base pointer callee saved register r31 is saved in the TOC slot of linkage area. This simplifies allocation of space for the base pointer and makes it convenient to locate programmatically and during debugging.
Dynamic Allocation¶
Note
TODO - More to come.
The NVPTX backend¶
The NVPTX code generator under lib/Target/NVPTX is an open-source version of the NVIDIA NVPTX code generator for LLVM. It is contributed by NVIDIA and is a port of the code generator used in the CUDA compiler (nvcc). It targets the PTX 3.0/3.1 ISA and can target any compute capability greater than or equal to 2.0 (Fermi).
This target is of production quality and should be completely compatible with the official NVIDIA toolchain.
Code Generator Options:
Option
Description
sm_20
Set shader model/compute capability to 2.0
sm_21
Set shader model/compute capability to 2.1
sm_30
Set shader model/compute capability to 3.0
sm_35
Set shader model/compute capability to 3.5
ptx30
Target PTX 3.0
ptx31
Target PTX 3.1
The extended Berkeley Packet Filter (eBPF) backend¶
Extended BPF (or eBPF) is similar to the original (“classic”) BPF (cBPF) used to filter network packets. The bpf() system call performs a range of operations related to eBPF. For both cBPF and eBPF programs, the Linux kernel statically analyzes the programs before loading them, in order to ensure that they cannot harm the running system. eBPF is a 64-bit RISC instruction set designed for one to one mapping to 64-bit CPUs. Opcodes are 8-bit encoded, and 87 instructions are defined. There are 10 registers, grouped by function as outlined below.
R0 return value from in-kernel functions; exit value for eBPF program
R1 - R5 function call arguments to in-kernel functions
R6 - R9 callee-saved registers preserved by in-kernel functions
R10 stack frame pointer (read only)
Instruction encoding (arithmetic and jump)¶
eBPF is reusing most of the opcode encoding from classic to simplify conversion of classic BPF to eBPF. For arithmetic and jump instructions the 8-bit ‘code’ field is divided into three parts:
+----------------+--------+--------------------+
| 4 bits | 1 bit | 3 bits |
| operation code | source | instruction class |
+----------------+--------+--------------------+
(MSB) (LSB)
Three LSB bits store instruction class which is one of:
BPF_LD 0x0
BPF_LDX 0x1
BPF_ST 0x2
BPF_STX 0x3
BPF_ALU 0x4
BPF_JMP 0x5
(unused) 0x6
BPF_ALU64 0x7
When BPF_CLASS(code) == BPF_ALU or BPF_ALU64 or BPF_JMP, 4th bit encodes source operand
BPF_X 0x1 use src_reg register as source operand
BPF_K 0x0 use 32 bit immediate as source operand
and four MSB bits store operation code
BPF_ADD 0x0 add
BPF_SUB 0x1 subtract
BPF_MUL 0x2 multiply
BPF_DIV 0x3 divide
BPF_OR 0x4 bitwise logical OR
BPF_AND 0x5 bitwise logical AND
BPF_LSH 0x6 left shift
BPF_RSH 0x7 right shift (zero extended)
BPF_NEG 0x8 arithmetic negation
BPF_MOD 0x9 modulo
BPF_XOR 0xa bitwise logical XOR
BPF_MOV 0xb move register to register
BPF_ARSH 0xc right shift (sign extended)
BPF_END 0xd endianness conversion
If BPF_CLASS(code) == BPF_JMP, BPF_OP(code) is one of
BPF_JA 0x0 unconditional jump
BPF_JEQ 0x1 jump ==
BPF_JGT 0x2 jump >
BPF_JGE 0x3 jump >=
BPF_JSET 0x4 jump if (DST & SRC)
BPF_JNE 0x5 jump !=
BPF_JSGT 0x6 jump signed >
BPF_JSGE 0x7 jump signed >=
BPF_CALL 0x8 function call
BPF_EXIT 0x9 function return
Instruction encoding (load, store)¶
For load and store instructions the 8-bit ‘code’ field is divided as:
+--------+--------+-------------------+
| 3 bits | 2 bits | 3 bits |
| mode | size | instruction class |
+--------+--------+-------------------+
(MSB) (LSB)
Size modifier is one of
BPF_W 0x0 word
BPF_H 0x1 half word
BPF_B 0x2 byte
BPF_DW 0x3 double word
Mode modifier is one of
BPF_IMM 0x0 immediate
BPF_ABS 0x1 used to access packet data
BPF_IND 0x2 used to access packet data
BPF_MEM 0x3 memory
(reserved) 0x4
(reserved) 0x5
BPF_XADD 0x6 exclusive add
Packet data access (BPF_ABS, BPF_IND)¶
Two non-generic instructions: (BPF_ABS | <size> | BPF_LD) and (BPF_IND | <size> | BPF_LD) which are used to access packet data. Register R6 is an implicit input that must contain pointer to sk_buff. Register R0 is an implicit output which contains the data fetched from the packet. Registers R1-R5 are scratch registers and must not be used to store the data across BPF_ABS | BPF_LD or BPF_IND | BPF_LD instructions. These instructions have implicit program exit condition as well. When eBPF program is trying to access the data beyond the packet boundary, the interpreter will abort the execution of the program.
- BPF_IND | BPF_W | BPF_LD is equivalent to:
R0 = ntohl(*(u32 *) (((struct sk_buff *) R6)->data + src_reg + imm32))
eBPF maps¶
eBPF maps are provided for sharing data between kernel and user-space. Currently implemented types are hash and array, with potential extension to support bloom filters, radix trees, etc. A map is defined by its type, maximum number of elements, key size and value size in bytes. eBPF syscall supports create, update, find and delete functions on maps.
Function calls¶
Function call arguments are passed using up to five registers (R1 - R5). The return value is passed in a dedicated register (R0). Four additional registers (R6 - R9) are callee-saved, and the values in these registers are preserved within kernel functions. R0 - R5 are scratch registers within kernel functions, and eBPF programs must therefor store/restore values in these registers if needed across function calls. The stack can be accessed using the read-only frame pointer R10. eBPF registers map 1:1 to hardware registers on x86_64 and other 64-bit architectures. For example, x86_64 in-kernel JIT maps them as
R0 - rax
R1 - rdi
R2 - rsi
R3 - rdx
R4 - rcx
R5 - r8
R6 - rbx
R7 - r13
R8 - r14
R9 - r15
R10 - rbp
since x86_64 ABI mandates rdi, rsi, rdx, rcx, r8, r9 for argument passing and rbx, r12 - r15 are callee saved.
Program start¶
An eBPF program receives a single argument and contains a single eBPF main routine; the program does not contain eBPF functions. Function calls are limited to a predefined set of kernel functions. The size of a program is limited to 4K instructions: this ensures fast termination and a limited number of kernel function calls. Prior to running an eBPF program, a verifier performs static analysis to prevent loops in the code and to ensure valid register usage and operand types.
The AMDGPU backend¶
The AMDGPU code generator lives in the lib/Target/AMDGPU
directory. This code generator is capable of targeting a variety of
AMD GPU processors. Refer to User Guide for AMDGPU Backend for more information.