llvm-project/llvm/lib/CodeGen/CodeGen.cpp

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//===-- CodeGen.cpp -------------------------------------------------------===//
//
// Part of the LLVM Project, under the Apache License v2.0 with LLVM Exceptions.
// See https://llvm.org/LICENSE.txt for license information.
// SPDX-License-Identifier: Apache-2.0 WITH LLVM-exception
//
//===----------------------------------------------------------------------===//
//
// This file implements the common initialization routines for the
// CodeGen library.
//
//===----------------------------------------------------------------------===//
#include "llvm-c/Initialization.h"
#include "llvm/InitializePasses.h"
#include "llvm/PassRegistry.h"
using namespace llvm;
/// initializeCodeGen - Initialize all passes linked into the CodeGen library.
void llvm::initializeCodeGen(PassRegistry &Registry) {
initializeAtomicExpandPass(Registry);
initializeBranchFolderPassPass(Registry);
initializeBranchRelaxationPass(Registry);
Correct dwarf unwind information in function epilogue This patch aims to provide correct dwarf unwind information in function epilogue for X86. It consists of two parts. The first part inserts CFI instructions that set appropriate cfa offset and cfa register in emitEpilogue() in X86FrameLowering. This part is X86 specific. The second part is platform independent and ensures that: * CFI instructions do not affect code generation (they are not counted as instructions when tail duplicating or tail merging) * Unwind information remains correct when a function is modified by different passes. This is done in a late pass by analyzing information about cfa offset and cfa register in BBs and inserting additional CFI directives where necessary. Added CFIInstrInserter pass: * analyzes each basic block to determine cfa offset and register are valid at its entry and exit * verifies that outgoing cfa offset and register of predecessor blocks match incoming values of their successors * inserts additional CFI directives at basic block beginning to correct the rule for calculating CFA Having CFI instructions in function epilogue can cause incorrect CFA calculation rule for some basic blocks. This can happen if, due to basic block reordering, or the existence of multiple epilogue blocks, some of the blocks have wrong cfa offset and register values set by the epilogue block above them. CFIInstrInserter is currently run only on X86, but can be used by any target that implements support for adding CFI instructions in epilogue. Patch by Violeta Vukobrat. Differential Revision: https://reviews.llvm.org/D42848 llvm-svn: 330706
2018-04-24 18:32:08 +08:00
initializeCFIInstrInserterPass(Registry);
initializeCodeGenPreparePass(Registry);
initializeDeadMachineInstructionElimPass(Registry);
initializeDetectDeadLanesPass(Registry);
initializeDwarfEHPreparePass(Registry);
initializeEarlyIfConverterPass(Registry);
initializeEarlyMachineLICMPass(Registry);
initializeEarlyTailDuplicatePass(Registry);
initializeExpandISelPseudosPass(Registry);
initializeExpandMemCmpPassPass(Registry);
initializeExpandPostRAPass(Registry);
initializeFEntryInserterPass(Registry);
initializeFinalizeMachineBundlesPass(Registry);
initializeFuncletLayoutPass(Registry);
initializeGCMachineCodeAnalysisPass(Registry);
initializeGCModuleInfoPass(Registry);
initializeIfConverterPass(Registry);
initializeImplicitNullChecksPass(Registry);
Introduce the "retpoline" x86 mitigation technique for variant #2 of the speculative execution vulnerabilities disclosed today, specifically identified by CVE-2017-5715, "Branch Target Injection", and is one of the two halves to Spectre.. Summary: First, we need to explain the core of the vulnerability. Note that this is a very incomplete description, please see the Project Zero blog post for details: https://googleprojectzero.blogspot.com/2018/01/reading-privileged-memory-with-side.html The basis for branch target injection is to direct speculative execution of the processor to some "gadget" of executable code by poisoning the prediction of indirect branches with the address of that gadget. The gadget in turn contains an operation that provides a side channel for reading data. Most commonly, this will look like a load of secret data followed by a branch on the loaded value and then a load of some predictable cache line. The attacker then uses timing of the processors cache to determine which direction the branch took *in the speculative execution*, and in turn what one bit of the loaded value was. Due to the nature of these timing side channels and the branch predictor on Intel processors, this allows an attacker to leak data only accessible to a privileged domain (like the kernel) back into an unprivileged domain. The goal is simple: avoid generating code which contains an indirect branch that could have its prediction poisoned by an attacker. In many cases, the compiler can simply use directed conditional branches and a small search tree. LLVM already has support for lowering switches in this way and the first step of this patch is to disable jump-table lowering of switches and introduce a pass to rewrite explicit indirectbr sequences into a switch over integers. However, there is no fully general alternative to indirect calls. We introduce a new construct we call a "retpoline" to implement indirect calls in a non-speculatable way. It can be thought of loosely as a trampoline for indirect calls which uses the RET instruction on x86. Further, we arrange for a specific call->ret sequence which ensures the processor predicts the return to go to a controlled, known location. The retpoline then "smashes" the return address pushed onto the stack by the call with the desired target of the original indirect call. The result is a predicted return to the next instruction after a call (which can be used to trap speculative execution within an infinite loop) and an actual indirect branch to an arbitrary address. On 64-bit x86 ABIs, this is especially easily done in the compiler by using a guaranteed scratch register to pass the target into this device. For 32-bit ABIs there isn't a guaranteed scratch register and so several different retpoline variants are introduced to use a scratch register if one is available in the calling convention and to otherwise use direct stack push/pop sequences to pass the target address. This "retpoline" mitigation is fully described in the following blog post: https://support.google.com/faqs/answer/7625886 We also support a target feature that disables emission of the retpoline thunk by the compiler to allow for custom thunks if users want them. These are particularly useful in environments like kernels that routinely do hot-patching on boot and want to hot-patch their thunk to different code sequences. They can write this custom thunk and use `-mretpoline-external-thunk` *in addition* to `-mretpoline`. In this case, on x86-64 thu thunk names must be: ``` __llvm_external_retpoline_r11 ``` or on 32-bit: ``` __llvm_external_retpoline_eax __llvm_external_retpoline_ecx __llvm_external_retpoline_edx __llvm_external_retpoline_push ``` And the target of the retpoline is passed in the named register, or in the case of the `push` suffix on the top of the stack via a `pushl` instruction. There is one other important source of indirect branches in x86 ELF binaries: the PLT. These patches also include support for LLD to generate PLT entries that perform a retpoline-style indirection. The only other indirect branches remaining that we are aware of are from precompiled runtimes (such as crt0.o and similar). The ones we have found are not really attackable, and so we have not focused on them here, but eventually these runtimes should also be replicated for retpoline-ed configurations for completeness. For kernels or other freestanding or fully static executables, the compiler switch `-mretpoline` is sufficient to fully mitigate this particular attack. For dynamic executables, you must compile *all* libraries with `-mretpoline` and additionally link the dynamic executable and all shared libraries with LLD and pass `-z retpolineplt` (or use similar functionality from some other linker). We strongly recommend also using `-z now` as non-lazy binding allows the retpoline-mitigated PLT to be substantially smaller. When manually apply similar transformations to `-mretpoline` to the Linux kernel we observed very small performance hits to applications running typical workloads, and relatively minor hits (approximately 2%) even for extremely syscall-heavy applications. This is largely due to the small number of indirect branches that occur in performance sensitive paths of the kernel. When using these patches on statically linked applications, especially C++ applications, you should expect to see a much more dramatic performance hit. For microbenchmarks that are switch, indirect-, or virtual-call heavy we have seen overheads ranging from 10% to 50%. However, real-world workloads exhibit substantially lower performance impact. Notably, techniques such as PGO and ThinLTO dramatically reduce the impact of hot indirect calls (by speculatively promoting them to direct calls) and allow optimized search trees to be used to lower switches. If you need to deploy these techniques in C++ applications, we *strongly* recommend that you ensure all hot call targets are statically linked (avoiding PLT indirection) and use both PGO and ThinLTO. Well tuned servers using all of these techniques saw 5% - 10% overhead from the use of retpoline. We will add detailed documentation covering these components in subsequent patches, but wanted to make the core functionality available as soon as possible. Happy for more code review, but we'd really like to get these patches landed and backported ASAP for obvious reasons. We're planning to backport this to both 6.0 and 5.0 release streams and get a 5.0 release with just this cherry picked ASAP for distros and vendors. This patch is the work of a number of people over the past month: Eric, Reid, Rui, and myself. I'm mailing it out as a single commit due to the time sensitive nature of landing this and the need to backport it. Huge thanks to everyone who helped out here, and everyone at Intel who helped out in discussions about how to craft this. Also, credit goes to Paul Turner (at Google, but not an LLVM contributor) for much of the underlying retpoline design. Reviewers: echristo, rnk, ruiu, craig.topper, DavidKreitzer Subscribers: sanjoy, emaste, mcrosier, mgorny, mehdi_amini, hiraditya, llvm-commits Differential Revision: https://reviews.llvm.org/D41723 llvm-svn: 323155
2018-01-23 06:05:25 +08:00
initializeIndirectBrExpandPassPass(Registry);
initializeInterleavedLoadCombinePass(Registry);
initializeInterleavedAccessPass(Registry);
initializeLiveDebugValuesPass(Registry);
initializeLiveDebugVariablesPass(Registry);
initializeLiveIntervalsPass(Registry);
initializeLiveRangeShrinkPass(Registry);
initializeLiveStacksPass(Registry);
initializeLiveVariablesPass(Registry);
initializeLocalStackSlotPassPass(Registry);
initializeLowerIntrinsicsPass(Registry);
initializeMIRCanonicalizerPass(Registry);
initializeMachineBlockFrequencyInfoPass(Registry);
Implement a block placement pass based on the branch probability and block frequency analyses. This differs substantially from the existing block-placement pass in LLVM: 1) It operates on the Machine-IR in the CodeGen layer. This exposes much more (and more precise) information and opportunities. Also, the results are more stable due to fewer transforms ocurring after the pass runs. 2) It uses the generalized probability and frequency analyses. These can model static heuristics, code annotation derived heuristics as well as eventual profile loading. By basing the optimization on the analysis interface it can work from any (or a combination) of these inputs. 3) It uses a more aggressive algorithm, both building chains from tho bottom up to maximize benefit, and using an SCC-based walk to layout chains of blocks in a profitable ordering without O(N^2) iterations which the old pass involves. The pass is currently gated behind a flag, and not enabled by default because it still needs to grow some important features. Most notably, it needs to support loop aligning and careful layout of loop structures much as done by hand currently in CodePlacementOpt. Once it supports these, and has sufficient testing and quality tuning, it should replace both of these passes. Thanks to Nick Lewycky and Richard Smith for help authoring & debugging this, and to Jakob, Andy, Eric, Jim, and probably a few others I'm forgetting for reviewing and answering all my questions. Writing a backend pass is *sooo* much better now than it used to be. =D llvm-svn: 142641
2011-10-21 14:46:38 +08:00
initializeMachineBlockPlacementPass(Registry);
initializeMachineBlockPlacementStatsPass(Registry);
initializeMachineCSEPass(Registry);
initializeMachineCombinerPass(Registry);
initializeMachineCopyPropagationPass(Registry);
initializeMachineDominatorTreePass(Registry);
initializeMachineFunctionPrinterPassPass(Registry);
initializeMachineLICMPass(Registry);
initializeMachineLoopInfoPass(Registry);
initializeMachineModuleInfoPass(Registry);
initializeMachineOptimizationRemarkEmitterPassPass(Registry);
initializeMachineOutlinerPass(Registry);
initializeMachinePipelinerPass(Registry);
initializeMachinePostDominatorTreePass(Registry);
initializeMachineRegionInfoPassPass(Registry);
initializeMachineSchedulerPass(Registry);
initializeMachineSinkingPass(Registry);
initializeMachineVerifierPassPass(Registry);
initializeOptimizePHIsPass(Registry);
initializePEIPass(Registry);
initializePHIEliminationPass(Registry);
initializePatchableFunctionPass(Registry);
initializePeepholeOptimizerPass(Registry);
initializePostMachineSchedulerPass(Registry);
initializePostRAHazardRecognizerPass(Registry);
[CodeGen] Add a new pass for PostRA sink Summary: This pass sinks COPY instructions into a successor block, if the COPY is not used in the current block and the COPY is live-in to a single successor (i.e., doesn't require the COPY to be duplicated). This avoids executing the the copy on paths where their results aren't needed. This also exposes additional opportunites for dead copy elimination and shrink wrapping. These copies were either not handled by or are inserted after the MachineSink pass. As an example of the former case, the MachineSink pass cannot sink COPY instructions with allocatable source registers; for AArch64 these type of copy instructions are frequently used to move function parameters (PhyReg) into virtual registers in the entry block.. For the machine IR below, this pass will sink %w19 in the entry into its successor (%bb.1) because %w19 is only live-in in %bb.1. ``` %bb.0: %wzr = SUBSWri %w1, 1 %w19 = COPY %w0 Bcc 11, %bb.2 %bb.1: Live Ins: %w19 BL @fun %w0 = ADDWrr %w0, %w19 RET %w0 %bb.2: %w0 = COPY %wzr RET %w0 ``` As we sink %w19 (CSR in AArch64) into %bb.1, the shrink-wrapping pass will be able to see %bb.0 as a candidate. With this change I observed 12% more shrink-wrapping candidate and 13% more dead copies deleted in spec2000/2006/2017 on AArch64. Reviewers: qcolombet, MatzeB, thegameg, mcrosier, gberry, hfinkel, john.brawn, twoh, RKSimon, sebpop, kparzysz Reviewed By: sebpop Subscribers: evandro, sebpop, sfertile, aemerson, mgorny, javed.absar, kristof.beyls, llvm-commits Differential Revision: https://reviews.llvm.org/D41463 llvm-svn: 328237
2018-03-23 04:06:47 +08:00
initializePostRAMachineSinkingPass(Registry);
initializePostRASchedulerPass(Registry);
initializePreISelIntrinsicLoweringLegacyPassPass(Registry);
initializeProcessImplicitDefsPass(Registry);
initializeRABasicPass(Registry);
initializeRAGreedyPass(Registry);
initializeRegAllocFastPass(Registry);
initializeRegUsageInfoCollectorPass(Registry);
initializeRegUsageInfoPropagationPass(Registry);
initializeRegisterCoalescerPass(Registry);
initializeRenameIndependentSubregsPass(Registry);
initializeSafeStackLegacyPassPass(Registry);
initializeScalarizeMaskedMemIntrinPass(Registry);
[ShrinkWrap] Add (a simplified version) of shrink-wrapping. This patch introduces a new pass that computes the safe point to insert the prologue and epilogue of the function. The interest is to find safe points that are cheaper than the entry and exits blocks. As an example and to avoid regressions to be introduce, this patch also implements the required bits to enable the shrink-wrapping pass for AArch64. ** Context ** Currently we insert the prologue and epilogue of the method/function in the entry and exits blocks. Although this is correct, we can do a better job when those are not immediately required and insert them at less frequently executed places. The job of the shrink-wrapping pass is to identify such places. ** Motivating example ** Let us consider the following function that perform a call only in one branch of a if: define i32 @f(i32 %a, i32 %b) { %tmp = alloca i32, align 4 %tmp2 = icmp slt i32 %a, %b br i1 %tmp2, label %true, label %false true: store i32 %a, i32* %tmp, align 4 %tmp4 = call i32 @doSomething(i32 0, i32* %tmp) br label %false false: %tmp.0 = phi i32 [ %tmp4, %true ], [ %a, %0 ] ret i32 %tmp.0 } On AArch64 this code generates (removing the cfi directives to ease readabilities): _f: ; @f ; BB#0: stp x29, x30, [sp, #-16]! mov x29, sp sub sp, sp, #16 ; =16 cmp w0, w1 b.ge LBB0_2 ; BB#1: ; %true stur w0, [x29, #-4] sub x1, x29, #4 ; =4 mov w0, wzr bl _doSomething LBB0_2: ; %false mov sp, x29 ldp x29, x30, [sp], #16 ret With shrink-wrapping we could generate: _f: ; @f ; BB#0: cmp w0, w1 b.ge LBB0_2 ; BB#1: ; %true stp x29, x30, [sp, #-16]! mov x29, sp sub sp, sp, #16 ; =16 stur w0, [x29, #-4] sub x1, x29, #4 ; =4 mov w0, wzr bl _doSomething add sp, x29, #16 ; =16 ldp x29, x30, [sp], #16 LBB0_2: ; %false ret Therefore, we would pay the overhead of setting up/destroying the frame only if we actually do the call. ** Proposed Solution ** This patch introduces a new machine pass that perform the shrink-wrapping analysis (See the comments at the beginning of ShrinkWrap.cpp for more details). It then stores the safe save and restore point into the MachineFrameInfo attached to the MachineFunction. This information is then used by the PrologEpilogInserter (PEI) to place the related code at the right place. This pass runs right before the PEI. Unlike the original paper of Chow from PLDI’88, this implementation of shrink-wrapping does not use expensive data-flow analysis and does not need hack to properly avoid frequently executed point. Instead, it relies on dominance and loop properties. The pass is off by default and each target can opt-in by setting the EnableShrinkWrap boolean to true in their derived class of TargetPassConfig. This setting can also be overwritten on the command line by using -enable-shrink-wrap. Before you try out the pass for your target, make sure you properly fix your emitProlog/emitEpilog/adjustForXXX method to cope with basic blocks that are not necessarily the entry block. ** Design Decisions ** 1. ShrinkWrap is its own pass right now. It could frankly be merged into PEI but for debugging and clarity I thought it was best to have its own file. 2. Right now, we only support one save point and one restore point. At some point we can expand this to several save point and restore point, the impacted component would then be: - The pass itself: New algorithm needed. - MachineFrameInfo: Hold a list or set of Save/Restore point instead of one pointer. - PEI: Should loop over the save point and restore point. Anyhow, at least for this first iteration, I do not believe this is interesting to support the complex cases. We should revisit that when we motivating examples. Differential Revision: http://reviews.llvm.org/D9210 <rdar://problem/3201744> llvm-svn: 236507
2015-05-06 01:38:16 +08:00
initializeShrinkWrapPass(Registry);
initializeSlotIndexesPass(Registry);
initializeStackColoringPass(Registry);
initializeStackMapLivenessPass(Registry);
initializeStackProtectorPass(Registry);
initializeStackSlotColoringPass(Registry);
initializeTailDuplicatePass(Registry);
initializeTargetPassConfigPass(Registry);
initializeTwoAddressInstructionPassPass(Registry);
initializeUnpackMachineBundlesPass(Registry);
initializeUnreachableBlockElimLegacyPassPass(Registry);
initializeUnreachableMachineBlockElimPass(Registry);
initializeVirtRegMapPass(Registry);
initializeVirtRegRewriterPass(Registry);
initializeWasmEHPreparePass(Registry);
initializeWinEHPreparePass(Registry);
initializeXRayInstrumentationPass(Registry);
}
void LLVMInitializeCodeGen(LLVMPassRegistryRef R) {
initializeCodeGen(*unwrap(R));
}