Spectre variant #1 for x86.
There is a lengthy, detailed RFC thread on llvm-dev which discusses the
high level issues. High level discussion is probably best there.
I've split the design document out of this patch and will land it
separately once I update it to reflect the latest edits and updates to
the Google doc used in the RFC thread.
This patch is really just an initial step. It isn't quite ready for
prime time and is only exposed via debugging flags. It has two major
limitations currently:
1) It only supports x86-64, and only certain ABIs. Many assumptions are
currently hard-coded and need to be factored out of the code here.
2) It doesn't include any options for more fine-grained control, either
of which control flow edges are significant or which loads are
important to be hardened.
3) The code is still quite rough and the testing lighter than I'd like.
However, this is enough for people to begin using. I have had numerous
requests from people to be able to experiment with this patch to
understand the trade-offs it presents and how to use it. We would also
like to encourage work to similar effect in other toolchains.
The ARM folks are actively developing a system based on this for
AArch64. We hope to merge this with their efforts when both are far
enough along. But we also don't want to block making this available on
that effort.
Many thanks to the *numerous* people who helped along the way here. For
this patch in particular, both Eric and Craig did a ton of review to
even have confidence in it as an early, rough cut at this functionality.
Differential Revision: https://reviews.llvm.org/D44824
llvm-svn: 336990
The key idea is to lower COPY nodes populating EFLAGS by scanning the
uses of EFLAGS and introducing dedicated code to preserve the necessary
state in a GPR. In the vast majority of cases, these uses are cmovCC and
jCC instructions. For such cases, we can very easily save and restore
the necessary information by simply inserting a setCC into a GPR where
the original flags are live, and then testing that GPR directly to feed
the cmov or conditional branch.
However, things are a bit more tricky if arithmetic is using the flags.
This patch handles the vast majority of cases that seem to come up in
practice: adc, adcx, adox, rcl, and rcr; all without taking advantage of
partially preserved EFLAGS as LLVM doesn't currently model that at all.
There are a large number of operations that techinaclly observe EFLAGS
currently but shouldn't in this case -- they typically are using DF.
Currently, they will not be handled by this approach. However, I have
never seen this issue come up in practice. It is already pretty rare to
have these patterns come up in practical code with LLVM. I had to resort
to writing MIR tests to cover most of the logic in this pass already.
I suspect even with its current amount of coverage of arithmetic users
of EFLAGS it will be a significant improvement over the current use of
pushf/popf. It will also produce substantially faster code in most of
the common patterns.
This patch also removes all of the old lowering for EFLAGS copies, and
the hack that forced us to use a frame pointer when EFLAGS copies were
found anywhere in a function so that the dynamic stack adjustment wasn't
a problem. None of this is needed as we now lower all of these copies
directly in MI and without require stack adjustments.
Lots of thanks to Reid who came up with several aspects of this
approach, and Craig who helped me work out a couple of things tripping
me up while working on this.
Differential Revision: https://reviews.llvm.org/D45146
llvm-svn: 329657
Summary:
The ShadowCallStack pass instruments functions marked with the
shadowcallstack attribute. The instrumented prolog saves the return
address to [gs:offset] where offset is stored and updated in [gs:0].
The instrumented epilog loads/updates the return address from [gs:0]
and checks that it matches the return address on the stack before
returning.
Reviewers: pcc, vitalybuka
Reviewed By: pcc
Subscribers: cryptoad, eugenis, craig.topper, mgorny, llvm-commits, kcc
Differential Revision: https://reviews.llvm.org/D44802
llvm-svn: 329139
If a load follows a store and reloads data that the store has written to memory, Intel microarchitectures can in many cases forward the data directly from the store to the load, This "store forwarding" saves cycles by enabling the load to directly obtain the data instead of accessing the data from cache or memory.
A "store forward block" occurs in cases that a store cannot be forwarded to the load. The most typical case of store forward block on Intel Core microarchiticutre that a small store cannot be forwarded to a large load.
The estimated penalty for a store forward block is ~13 cycles.
This pass tries to recognize and handle cases where "store forward block" is created by the compiler when lowering memcpy calls to a sequence
of a load and a store.
The pass currently only handles cases where memcpy is lowered to XMM/YMM registers, it tries to break the memcpy into smaller copies.
breaking the memcpy should be possible since there is no atomicity guarantee for loads and stores to XMM/YMM.
Differential revision: https://reviews.llvm.org/D41330
Change-Id: Ib48836ccdf6005989f7d4466fa2035b7b04415d9
llvm-svn: 328973
If a load follows a store and reloads data that the store has written to memory, Intel microarchitectures can in many cases forward the data directly from the store to the load, This "store forwarding" saves cycles by enabling the load to directly obtain the data instead of accessing the data from cache or memory.
A "store forward block" occurs in cases that a store cannot be forwarded to the load. The most typical case of store forward block on Intel Core microarchiticutre that a small store cannot be forwarded to a large load.
The estimated penalty for a store forward block is ~13 cycles.
This pass tries to recognize and handle cases where "store forward block" is created by the compiler when lowering memcpy calls to a sequence
of a load and a store.
The pass currently only handles cases where memcpy is lowered to XMM/YMM registers, it tries to break the memcpy into smaller copies.
breaking the memcpy should be possible since there is no atomicity guarantee for loads and stores to XMM/YMM.
Change-Id: Ic41aa9ade6512e0478db66e07e2fde41b4fb35f9
llvm-svn: 325128
It asserts building Chromium; see PR36346.
(This also reverts the follow-up r324836.)
> If a load follows a store and reloads data that the store has written to memory, Intel microarchitectures can in many cases forward the data directly from the store to the load, This "store forwarding" saves cycles by enabling the load to directly obtain the data instead of accessing the data from cache or memory.
> A "store forward block" occurs in cases that a store cannot be forwarded to the load. The most typical case of store forward block on Intel Core microarchiticutre that a small store cannot be forwarded to a large load.
> The estimated penalty for a store forward block is ~13 cycles.
>
> This pass tries to recognize and handle cases where "store forward block" is created by the compiler when lowering memcpy calls to a sequence
> of a load and a store.
>
> The pass currently only handles cases where memcpy is lowered to XMM/YMM registers, it tries to break the memcpy into smaller copies.
> breaking the memcpy should be possible since there is no atomicity guarantee for loads and stores to XMM/YMM.
llvm-svn: 324887
If a load follows a store and reloads data that the store has written to memory, Intel microarchitectures can in many cases forward the data directly from the store to the load, This "store forwarding" saves cycles by enabling the load to directly obtain the data instead of accessing the data from cache or memory.
A "store forward block" occurs in cases that a store cannot be forwarded to the load. The most typical case of store forward block on Intel Core microarchiticutre that a small store cannot be forwarded to a large load.
The estimated penalty for a store forward block is ~13 cycles.
This pass tries to recognize and handle cases where "store forward block" is created by the compiler when lowering memcpy calls to a sequence
of a load and a store.
The pass currently only handles cases where memcpy is lowered to XMM/YMM registers, it tries to break the memcpy into smaller copies.
breaking the memcpy should be possible since there is no atomicity guarantee for loads and stores to XMM/YMM.
Change-Id: I620b6dc91583ad9a1444591e3ddc00dd25d81748
llvm-svn: 324835
Summary:
This removes the need for a machine module pass using some deeply
questionable hacks. This should address PR36123 which is a case where in
full LTO the memory usage of a machine module pass actually ended up
being significant.
We should revert this on trunk as soon as we understand and fix the
memory usage issue, but we should include this in any backports of
retpolines themselves.
Reviewers: echristo, MatzeB
Subscribers: sanjoy, mcrosier, mehdi_amini, hiraditya, llvm-commits
Differential Revision: https://reviews.llvm.org/D42726
llvm-svn: 323915
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
CET (Control-Flow Enforcement Technology) introduces a new mechanism called IBT (Indirect Branch Tracking).
According to IBT, each Indirect branch should land on dedicated ENDBR instruction (End Branch).
The new pass adds ENDBR instructions for every indirect jmp/call (including jumps using jump tables / switches).
For more information, please see the following:
https://software.intel.com/sites/default/files/managed/4d/2a/control-flow-enforcement-technology-preview.pdf
Differential Revision: https://reviews.llvm.org/D40482
Change-Id: Icb754489faf483a95248f96982a4e8b1009eb709
llvm-svn: 322062
The pass scans the function to find instruction chains that define
registers in the same domain (closures).
It then calculates the cost of converting the closure to another domain.
If found profitable, the instructions are converted to instructions in
the other domain and the register classes are changed accordingly.
This commit adds the pass infrastructure and a simple conversion from
the GPR domain to the Mask domain.
Differential Revision:
https://reviews.llvm.org/D37251
Change-Id: Ic2cf1d76598110401168326d411128ae2580a604
llvm-svn: 316288
LLVM compiler recognizes opportunities to transform a branch into IR select instruction(s) - later it will be lowered into X86::CMOV instruction, assuming no other optimization eliminated the SelectInst.
However, it is not always profitable to emit X86::CMOV instruction. For example, branch is preferable over an X86::CMOV instruction when:
1. Branch is well predicted
2. Condition operand is expensive, compared to True-value and the False-value operands
In CodeGenPrepare pass there is a shallow optimization that tries to convert SelectInst into branch, but it is not enough.
This commit, implements machine optimization pass that converts X86::CMOV instruction(s) into branch, based on a conservative heuristic.
Differential Revision: https://reviews.llvm.org/D34769
llvm-svn: 308142
Summary:
The SelectionDAG importer now imports rules with Predicate's attached via
Requires, PredicateControl, etc. These predicates are implemented as
bitset's to allow multiple predicates to be tested together. However,
unlike the MC layer subtarget features, each target only pays for it's own
predicates (e.g. AArch64 doesn't have 192 feature bits just because X86
needs a lot).
Both AArch64 and X86 derive at least one predicate from the MachineFunction
or Function so they must re-initialize AvailableFeatures before each
function. They also declare locals in <Target>InstructionSelector so that
computeAvailableFeatures() can use the code from SelectionDAG without
modification.
Reviewers: rovka, qcolombet, aditya_nandakumar, t.p.northover, ab
Reviewed By: rovka
Subscribers: aemerson, rengolin, dberris, kristof.beyls, llvm-commits, igorb
Differential Revision: https://reviews.llvm.org/D31418
llvm-svn: 300993
It's causing llvm-clang-x86_64-expensive-checks-win to fail to compile and I
haven't worked out why. Reverting to make it green while I figure it out.
llvm-svn: 300978
Summary:
The SelectionDAG importer now imports rules with Predicate's attached via
Requires, PredicateControl, etc. These predicates are implemented as
bitset's to allow multiple predicates to be tested together. However,
unlike the MC layer subtarget features, each target only pays for it's own
predicates (e.g. AArch64 doesn't have 192 feature bits just because X86
needs a lot).
Both AArch64 and X86 derive at least one predicate from the MachineFunction
or Function so they must re-initialize AvailableFeatures before each
function. They also declare locals in <Target>InstructionSelector so that
computeAvailableFeatures() can use the code from SelectionDAG without
modification.
Reviewers: rovka, qcolombet, aditya_nandakumar, t.p.northover, ab
Reviewed By: rovka
Subscribers: aemerson, rengolin, dberris, kristof.beyls, llvm-commits, igorb
Differential Revision: https://reviews.llvm.org/D31418
llvm-svn: 300964
Summary: This resolves the issue of tablegen-erated includes in the headers for non-GlobalISel builds in a simpler way than before.
Reviewers: qcolombet, ab
Reviewed By: ab
Subscribers: igorb, ab, mgorny, dberris, rovka, llvm-commits, kristof.beyls
Differential Revision: https://reviews.llvm.org/D30998
llvm-svn: 299637
There are cases of AVX-512 instructions that have two possible encodings. This is the case with instructions that use vector registers with low indexes of 0 - 15 and do not use the zmm registers or the mask k registers.
The EVEX encoding prefix requires 4 bytes whereas the VEX prefix can take only up to 3 bytes. Consequently, using the VEX encoding for these instructions results in a code size reduction of ~2 bytes even though it is compiled with the AVX-512 features enabled.
Reviewers: Craig Topper, Zvi Rackoover, Elena Demikhovsky
Differential Revision: https://reviews.llvm.org/D27901
llvm-svn: 290663
xorl + setcc is generally the preferred sequence due to the partial register
stall setcc + movzbl suffers from. As a bonus, it also encodes one byte smaller.
This fixes PR28146.
The original commit tried inserting an 8bit-subreg into a GR32 (not GR32_ABCD)
which was not appreciated by fast regalloc on 32-bit.
llvm-svn: 274802
xorl + setcc is generally the preferred sequence due to the partial register
stall setcc + movzbl suffers from. As a bonus, it also encodes one byte smaller.
This fixes PR28146.
Differential Revision: http://reviews.llvm.org/D21774
llvm-svn: 274692
with an additional fix to make RegAllocFast ignore undef physreg uses. It would
previously get confused about the "push %eax" instruction's use of eax. That
method for adjusting the stack pointer is used in X86FrameLowering::emitSPUpdate
as well, but since that runs after register-allocation, we didn't run into the
RegAllocFast issue before.
llvm-svn: 269949
This patch moves the expansion of WIN_ALLOCA pseudo-instructions
into a separate pass that walks the CFG and lowers the instructions
based on a conservative estimate of the offset between the stack
pointer and the lowest accessed stack address.
The goal is to reduce binary size and run-time costs by removing
calls to _chkstk. While it doesn't fix all the code quality problems
with inalloca calls, it's an incremental improvement for PR27076.
Differential Revision: http://reviews.llvm.org/D20263
llvm-svn: 269828
Make x86 OptimizeLEAs pass remove LEA instruction if there is another LEA
(in the same basic block) which calculates address differing only be a
displacement. Works only for -Oz.
Differential Revision: http://reviews.llvm.org/D13295
llvm-svn: 257589
Add new x86 pass which replaces address calculations in load or store instructions with def register of existing LEA (must be in the same basic block), if the LEA calculates address that differs only by a displacement. Works only with -Os or -Oz.
Differential Revision: http://reviews.llvm.org/D13294
llvm-svn: 254712
The patch is generated using this command:
tools/clang/tools/extra/clang-tidy/tool/run-clang-tidy.py -fix \
-checks=-*,llvm-namespace-comment -header-filter='llvm/.*|clang/.*' \
llvm/lib/
Thanks to Eugene Kosov for the original patch!
llvm-svn: 240137
The problem was that I slipped a change required for shrink-wrapping, namely I
used getFirstTerminator instead of the getLastNonDebugInstr that was here before
the refactoring, whereas the surrounding code is not yet patched for that.
Original message:
[X86] Refactor the prologue emission to prepare for shrink-wrapping.
- Add a late pass to expand pseudo instructions (tail call and EH returns).
Instead of doing it in the prologue emission.
- Factor some static methods in X86FrameLowering to ease code sharing.
NFC.
Related to <rdar://problem/20821487>
llvm-svn: 238035
Revert "[X86] Refactor the prologue emission to prepare for shrink-wrapping."
This reverts commit 6b3b93fc8b68a2c806aa992ee4bd3d7f61898d4b.
This reverts commit ab0b15dff8539826283a59c2dd700a18a9680e0f.
llvm-svn: 238011
- Add a late pass to expand pseudo instructions (tail call and EH returns).
Instead of doing it in the prologue emission.
- Factor some static methods in X86FrameLowering to ease code sharing.
NFC.
Related to <rdar://problem/20821487>
llvm-svn: 237977
This reverts commit r236360.
This change exposed a bug in WinEHPrepare by opting win32 code into EH
preparation. We already knew that WinEHPrepare has bugs, and is the
status quo for x64, so I don't think that's a reason to hold off on this
change. I disabled exceptions in the sanitizer tests in r236505 and an
earlier revision.
llvm-svn: 236508
This pass is responsible for constructing the EH registration object
that gets linked into fs:00, which is all it does in this change. In the
future, it will also insert stores to update the EH state number.
I considered keeping this functionality in WinEHPrepare, but it's pretty
separable and X86 specific. It has conceptually very little to do with
the task of WinEHPrepare, which is currently outlining. WinEHPrepare is
also in theory useful on ARM, but this logic is pretty x86 specific.
Reviewers: andrew.w.kaylor, majnemer
Differential Revision: http://reviews.llvm.org/D9422
llvm-svn: 236339
This moves the transformation introduced in r223757 into a separate MI pass.
This allows it to cover many more cases (not only cases where there must be a
reserved call frame), and perform rudimentary call folding. It still doesn't
have a heuristic, so it is enabled only for optsize/minsize, with stack
alignment <= 8, where it ought to be a fairly clear win.
(Re-commit of r227728)
Differential Revision: http://reviews.llvm.org/D6789
llvm-svn: 227752
This moves the transformation introduced in r223757 into a separate MI pass.
This allows it to cover many more cases (not only cases where there must be a
reserved call frame), and perform rudimentary call folding. It still doesn't
have a heuristic, so it is enabled only for optsize/minsize, with stack
alignment <= 8, where it ought to be a fairly clear win.
Differential Revision: http://reviews.llvm.org/D6789
llvm-svn: 227728
This required a new hook called hasLoadLinkedStoreConditional to know whether
to expand atomics to LL/SC (ARM, AArch64, in a future patch Power) or to
CmpXchg (X86).
Apart from that, the new code in AtomicExpandPass is mostly moved from
X86AtomicExpandPass. The main result of this patch is to get rid of that
pass, which had lots of code duplicated with AtomicExpandPass.
llvm-svn: 217928
Chandler Carruth