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Add check_mem_access(addr) function to check if memory at addr can
be written or read returning false if memory is not accessible.
This function changes signal handler for SIGSEGV and SIGBUS signals
when it is called first, and it is not thread-safe.
Co-authored-by: Adhemerval Zanella Netto <adhemerval.zanella@linaro.org>
Reviewed-by: Adhemerval Zanella <adhemerval.zanella@linaro.org>
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Both NPTL and HTL support PTHREAD_PROCESS_SHARED, and since the removal of
the NaCL port there are no other pthread implementations.
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The symbol was unintentionally leaked on ports introduced after
GLIBC_2.34, provide the compat symbol to avoid breaking ABI on them.
Signed-off-by: Xi Ruoyao <xry111@xry111.site>
Reviewed-by: Florian Weimer <fweimer@redhat.com>
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This patch replaces _dl_stack_flags global variable by
_dl_stack_prot_flags.
The advantage is that any convertion from p_flags to final used mprotect
flags occurs at loading of p_flags. It avoids repeated spurious
convertions of _dl_stack_flags, for example in allocate_thread_stack.
This modification was suggested in:
https://sourceware.org/pipermail/libc-alpha/2025-March/165537.html
Reviewed-by: Adhemerval Zanella <adhemerval.zanella@linaro.org>
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The SYSCALL_CANCEL calls __syscall_cancel, which in turn
calls __internal_syscall_cancel with an 'int' return instead of the
expected 'long int'. This causes issues with syscalls that return
values larger than INT_MAX, such as copy_file_range [1].
Checked on x86_64-linux-gnu.
[1] https://debbugs.gnu.org/cgi/bugreport.cgi?bug=79139
Reviewed-by: Andreas K. Huettel <dilfridge@gentoo.org>
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The SIGCANCEL signal handler should not issue __syscall_do_cancel,
which calls __do_cancel and __pthread_unwind, if the cancellation
is already in proces (and libgcc unwind is not reentrant). Any
cancellation signal received after is ignored.
Checked on x86_64-linux-gnu and aarch64-linux-gnu.
Tested-by: Aurelien Jarno <aurelien@aurel32.net>
Reviewed-by: Florian Weimer <fweimer@redhat.com>
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Do not add the pthread_atfork routine again in nptl/Makefile,
instead rely on sysdeps/pthread/Makefile for the integration
(as this is the directory that contains the source file).
In sysdeps/pthread/Makefile, add to static-only-routines.
Reviewed-by: Joseph Myers <josmyers@redhat.com>
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The test uses the while (true) construct.
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Current Bionic has this function, with enhanced error checking
(the undefined case terminates the process).
Reviewed-by: Joseph Myers <josmyers@redhat.com>
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When GNU Binutils is configured with --enable-error-execstack=yes, a handful
of our tests which rely on -Wl,-z,execstack fail. Pass --Wl,--no-error-execstack
to override the behaviour and get a warning instead.
Bug: https://sourceware.org/PR32717
Reviewed-by: Adhemerval Zanella <adhemerval.zanella@linaro.org>
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In the block comment for __pthread_cond_wait_common we mention
__g_refs, but the implementation no longer uses group references.
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Decorate BSS mappings with [anon: glibc: .bss <file>], for example
[anon: glibc: .bss /lib/libc.so.6]. The string ".bss" is already used
by bionic so use the same, but add the filename as well. If the name
would be longer than what the kernel allows, drop the directory part
of the path.
Refactor glibc.mem.decorate_maps check to a separate function and use
it to avoid assembling a name, which would not be used later.
Signed-off-by: Petr Malat <oss@malat.biz>
Reviewed-by: Adhemerval Zanella <adhemerval.zanella@linaro.org>
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Linux 6.13 (662df3e5c3766) added a lightweight way to define guard areas
through madvise syscall. Instead of PROT_NONE the guard region through
mprotect, userland can madvise the same area with a special flag, and
the kernel ensures that accessing the area will trigger a SIGSEGV (as for
PROT_NONE mapping).
The madvise way has the advantage of less kernel memory consumption for
the process page-table (one less VMA per guard area), and slightly less
contention on kernel (also due to the fewer VMA areas being tracked).
The pthread_create allocates a new thread stack in two ways: if a guard
area is set (the default) it allocates the memory range required using
PROT_NONE and then mprotect the usable stack area. Otherwise, if a
guard page is not set it allocates the region with the required flags.
For the MADV_GUARD_INSTALL support, the stack area region is allocated
with required flags and then the guard region is installed. If the
kernel does not support it, the usual way is used instead (and
MADV_GUARD_INSTALL is disabled for future stack creations).
The stack allocation strategy is recorded on the pthread struct, and it
is used in case the guard region needs to be resized. To avoid needing
an extra field, the 'user_stack' is repurposed and renamed to 'stack_mode'.
This patch also adds a proper test for the pthread guard.
I checked on x86_64, aarch64, powerpc64le, and hppa with kernel 6.13.0-rc7.
Reviewed-by: DJ Delorie <dj@redhat.com>
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Set stack size attribute to the size of the mmap'd region only
when the size of the remaining stack space is less than the size
of the mmap'd region.
This was reversed. As a result, the initial stack size was only
135168 bytes. On architectures where the stack grows down, the
initial stack size is approximately 8384512 bytes with the default
rlimit settings. The small main stack size on hppa broke
applications like ruby that check for stack overflows.
Signed-off-by: John David Anglin <dave.anglin@bell.net>
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The file uses the identifiers bool, false, true.
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The LSB of g_signals was unused. The LSB of g1_start was used to indicate
which group is G2. This was used to always go to sleep in pthread_cond_wait
if a waiter is in G2. A comment earlier in the file says that this is not
correct to do:
"Waiters cannot determine whether they are currently in G2 or G1 -- but they
do not have to because all they are interested in is whether there are
available signals"
I either would have had to update the comment, or get rid of the check. I
chose to get rid of the check. In fact I don't quite know why it was there.
There will never be available signals for group G2, so we didn't need the
special case. Even if there were, this would just be a spurious wake. This
might have caught some cases where the count has wrapped around, but it
wouldn't reliably do that, (and even if it did, why would you want to force a
sleep in that case?) and we don't support that many concurrent waiters
anyway. Getting rid of it allows us to use one more bit, making us more
robust to wraparound.
Signed-off-by: Malte Skarupke <malteskarupke@fastmail.fm>
Reviewed-by: Carlos O'Donell <carlos@redhat.com>
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This function no longer waits for threads to leave g1, so rename it to
__condvar_switch_g1
Signed-off-by: Malte Skarupke <malteskarupke@fastmail.fm>
Reviewed-by: Carlos O'Donell <carlos@redhat.com>
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In my previous change I turned a nested loop into a simple loop. I'm doing
the resulting indentation changes in a separate commit to make the diff on
the previous commit easier to review.
Signed-off-by: Malte Skarupke <malteskarupke@fastmail.fm>
Reviewed-by: Carlos O'Donell <carlos@redhat.com>
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The loop was a little more complicated than necessary. There was only one
break statement out of the inner loop, and the outer loop was nearly empty.
So just remove the outer loop, moving its code to the one break statement in
the inner loop. This allows us to replace all gotos with break statements.
Signed-off-by: Malte Skarupke <malteskarupke@fastmail.fm>
Reviewed-by: Carlos O'Donell <carlos@redhat.com>
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This variable used to be needed to wait in group switching until all sleepers
have confirmed that they have woken. This is no longer needed. Nothing waits
on this variable so there is no need to track how many threads are currently
asleep in each group.
Signed-off-by: Malte Skarupke <malteskarupke@fastmail.fm>
Reviewed-by: Carlos O'Donell <carlos@redhat.com>
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pthread_cond_wait was checking whether it was in a closed group no less than
four times. Checking once is enough. Here are the four checks:
1. While spin-waiting. This was dead code: maxspin is set to 0 and has been
for years.
2. Before deciding to go to sleep, and before incrementing grefs: I kept this
3. After incrementing grefs. There is no reason to think that the group would
close while we do an atomic increment. Obviously it could close at any
point, but that doesn't mean we have to recheck after every step. This
check was equally good as check 2, except it has to do more work.
4. When we find ourselves in a group that has a signal. We only get here after
we check that we're not in a closed group. There is no need to check again.
The check would only have helped in cases where the compare_exchange in the
next line would also have failed. Relying on the compare_exchange is fine.
Removing the duplicate checks clarifies the code.
Signed-off-by: Malte Skarupke <malteskarupke@fastmail.fm>
Reviewed-by: Carlos O'Donell <carlos@redhat.com>
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This wake is unnecessary. We only switch groups after every sleeper in a group
has been woken. Sure, they may take a while to actually wake up and may still
hold a reference, but waking them a second time doesn't speed that up. Instead
this just makes the code more complicated and may hide problems.
In particular this safety wake wouldn't even have helped with the bug that was
fixed by Barrus' patch: The bug there was that pthread_cond_signal would not
switch g1 when it should, so we wouldn't even have entered this code path.
Signed-off-by: Malte Skarupke <malteskarupke@fastmail.fm>
Reviewed-by: Carlos O'Donell <carlos@redhat.com>
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Some comments were wrong after the most recent commit. This fixes that.
Also fixing indentation where it was using spaces instead of tabs.
Signed-off-by: Malte Skarupke <malteskarupke@fastmail.fm>
Reviewed-by: Carlos O'Donell <carlos@redhat.com>
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This fixes the lost wakeup (from a bug in signal stealing) with a change
in the usage of g_signals[] in the condition variable internal state.
It also completely eliminates the concept and handling of signal stealing,
as well as the need for signalers to block to wait for waiters to wake
up every time there is a G1/G2 switch. This greatly reduces the average
and maximum latency for pthread_cond_signal.
The g_signals[] field now contains a signal count that is relative to
the current g1_start value. Since it is a 32-bit field, and the LSB is
still reserved (though not currently used anymore), it has a 31-bit value
that corresponds to the low 31 bits of the sequence number in g1_start.
(since g1_start also has an LSB flag, this means bits 31:1 in g_signals
correspond to bits 31:1 in g1_start, plus the current signal count)
By making the signal count relative to g1_start, there is no longer
any ambiguity or A/B/A issue, and thus any checks before blocking,
including the futex call itself, are guaranteed not to block if the G1/G2
switch occurs, even if the signal count remains the same. This allows
initially safely blocking in G2 until the switch to G1 occurs, and
then transitioning from G1 to a new G1 or G2, and always being able to
distinguish the state change. This removes the race condition and A/B/A
problems that otherwise ocurred if a late (pre-empted) waiter were to
resume just as the futex call attempted to block on g_signal since
otherwise there was no last opportunity to re-check things like whether
the current G1 group was already closed.
By fixing these issues, the signal stealing code can be eliminated,
since there is no concept of signal stealing anymore. The code to block
for all waiters to exit g_refs can also be removed, since any waiters
that are still in the g_refs region can be guaranteed to safely wake
up and exit. If there are still any left at this time, they are all
sent one final futex wakeup to ensure that they are not blocked any
longer, but there is no need for the signaller to block and wait for
them to wake up and exit the g_refs region.
The signal count is then effectively "zeroed" but since it is now
relative to g1_start, this is done by advancing it to a new value that
can be observed by any pending blocking waiters. Any late waiters can
always tell the difference, and can thus just cleanly exit if they are
in a stale G1 or G2. They can never steal a signal from the current
G1 if they are not in the current G1, since the signal value that has
to match in the cmpxchg has the low 31 bits of the g1_start value
contained in it, and that's first checked, and then it won't match if
there's a G1/G2 change.
Note: the 31-bit sequence number used in g_signals is designed to
handle wrap-around when checking the signal count, but if the entire
31-bit wraparound (2 billion signals) occurs while there is still a
late waiter that has not yet resumed, and it happens to then match
the current g1_start low bits, and the pre-emption occurs after the
normal "closed group" checks (which are 64-bit) but then hits the
futex syscall and signal consuming code, then an A/B/A issue could
still result and cause an incorrect assumption about whether it
should block. This particular scenario seems unlikely in practice.
Note that once awake from the futex, the waiter would notice the
closed group before consuming the signal (since that's still a 64-bit
check that would not be aliased in the wrap-around in g_signals),
so the biggest impact would be blocking on the futex until the next
full wakeup from a G1/G2 switch.
Signed-off-by: Frank Barrus <frankbarrus_sw@shaggy.cc>
Reviewed-by: Carlos O'Donell <carlos@redhat.com>
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Some kernels on S390 appear to return a CPU affinity mask based on
configured processors rather than the ones online. Overallocate the CPU
set to match that, but operate only on the ones online.
Signed-off-by: Siddhesh Poyarekar <siddhesh@sourceware.org>
Co-authored-by: Siddhesh Poyarekar <siddhesh@sourceware.org>
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The rseq extensible ABI implementation moved the rseq area to the 'extra
TLS' block, remove the unused 'rseq_area' member of 'struct pthread'.
Signed-off-by: Michael Jeanson <mjeanson@efficios.com>
Reviewed-by: Mathieu Desnoyers <mathieu.desnoyers@efficios.com>
Reviewed-by: Florian Weimer <fweimer@redhat.com>
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Move the rseq area to the newly added 'extra TLS' block, this is the
last step in adding support for the rseq extended ABI. The size of the
rseq area is now dynamic and depends on the rseq features reported by
the kernel through the elf auxiliary vector. This will allow
applications to use rseq features past the 32 bytes of the original rseq
ABI as they become available in future kernels.
Signed-off-by: Michael Jeanson <mjeanson@efficios.com>
Reviewed-by: Mathieu Desnoyers <mathieu.desnoyers@efficios.com>
Reviewed-by: Florian Weimer <fweimer@redhat.com>
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Add a couple of tests to verify that CPU affinity set using
sched_setaffinity and pthread_setaffinity_np are inherited by a child
process and child thread.
Signed-off-by: Siddhesh Poyarekar <siddhesh@sourceware.org>
Reviewed-by: Adhemerval Zanella <adhemerval.zanella@linaro.org>
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This reverts commit 7c22dcda27743658b6b8ea479283b384ad56bd5a.
The padding is required by Chromium's MaybeUpdateGlibcTidCache
in sandbox/linux/services/namespace_sandbox.cc.
Reviewed-by: Adhemerval Zanella <adhemerval.zanella@linaro.org>
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If some shared library loaded with dlopen/dlmopen requires an executable
stack, either implicitly because of a missing GNU_STACK ELF header
(where the ABI default flags implies in the executable bit) or explicitly
because of the executable bit from GNU_STACK; the loader will try to set
the both the main thread and all thread stacks (from the pthread cache)
as executable.
Besides the issue where any __nptl_change_stack_perm failure does not
undo the previous executable transition (meaning that if the library
fails to load, there can be thread stacks with executable stacks), this
behavior was used on a CVE [1] as a vector for RCE.
This patch changes that if a shared library requires an executable
stack, and the current stack is not executable, dlopen fails. The
change is done only for dynamically loaded modules, if the program
or any dependency requires an executable stack, the loader will still
change the main thread before program execution and any thread created
with default stack configuration.
[1] https://www.qualys.com/2023/07/19/cve-2023-38408/rce-openssh-forwarded-ssh-agent.txt
Checked on x86_64-linux-gnu and i686-linux-gnu.
Reviewed-by: Florian Weimer <fweimer@redhat.com>
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The previous use of padding within a union made it impossible to
re-use the padding for GLIBC_PRIVATE ABI preservation because
tcbhead_t could use up all of the padding (as was historically the
case on x86-64). Allocating padding unconditionally addresses this
issue.
Reviewed-by: Adhemerval Zanella <adhemerval.zanella@linaro.org>
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Add __attribute_optimization_barrier__ to disable inlining and cloning on a
function. For Clang, expand it to
__attribute__ ((optnone))
Otherwise, expand it to
__attribute__ ((noinline, clone))
Co-Authored-By: H.J. Lu <hjl.tools@gmail.com>
Reviewed-by: Sam James <sam@gentoo.org>
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Clang doesn't support -Wmaybe-uninitialized.
Signed-off-by: H.J. Lu <hjl.tools@gmail.com>
Reviewed-by: Sam James <sam@gentoo.org>
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Since trampoline is required to test execstack, enable execstack tests
only if compiler supports trampoline.
Signed-off-by: H.J. Lu <hjl.tools@gmail.com>
Reviewed-by: Sam James <sam@gentoo.org>
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Add a descriptive comment to the tst-pthread-cpuclockid-invalid test and
also drop pthread_getcpuclockid from the TODO-testing list since it now
has full coverage.
Signed-off-by: Siddhesh Poyarekar <siddhesh@sourceware.org>
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Replace 0 by NULL and {0} by {}.
Omit a few cases that aren't so trivial to fix.
Link: <https://gcc.gnu.org/bugzilla/show_bug.cgi?id=117059>
Link: <https://software.codidact.com/posts/292718/292759#answer-292759>
Signed-off-by: Alejandro Colomar <alx@kernel.org>
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Exercise the case where an exited thread will cause
pthread_getcpuclockid to fail.
Signed-off-by: Siddhesh Poyarekar <siddhesh@sourceware.org>
Reviewed-by: Florian Weimer <fweimer@redhat.com>
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Linux 6.11 has getrandom() in vDSO. It operates on a thread-local opaque
state allocated with mmap using flags specified by the vDSO.
Multiple states are allocated at once, as many as fit into a page, and
these are held in an array of available states to be doled out to each
thread upon first use, and recycled when a thread terminates. As these
states run low, more are allocated.
To make this procedure async-signal-safe, a simple guard is used in the
LSB of the opaque state address, falling back to the syscall if there's
reentrancy contention.
Also, _Fork() is handled by blocking signals on opaque state allocation
(so _Fork() always sees a consistent state even if it interrupts a
getrandom() call) and by iterating over the thread stack cache on
reclaim_stack. Each opaque state will be in the free states list
(grnd_alloc.states) or allocated to a running thread.
The cancellation is handled by always using GRND_NONBLOCK flags while
calling the vDSO, and falling back to the cancellable syscall if the
kernel returns EAGAIN (would block). Since getrandom is not defined by
POSIX and cancellation is supported as an extension, the cancellation is
handled as 'may occur' instead of 'shall occur' [1], meaning that if
vDSO does not block (the expected behavior) getrandom will not act as a
cancellation entrypoint. It avoids a pthread_testcancel call on the fast
path (different than 'shall occur' functions, like sem_wait()).
It is currently enabled for x86_64, which is available in Linux 6.11,
and aarch64, powerpc32, powerpc64, loongarch64, and s390x, which are
available in Linux 6.12.
Link: https://pubs.opengroup.org/onlinepubs/9799919799/nframe.html [1]
Co-developed-by: Jason A. Donenfeld <Jason@zx2c4.com>
Tested-by: Jason A. Donenfeld <Jason@zx2c4.com> # x86_64
Tested-by: Adhemerval Zanella <adhemerval.zanella@linaro.org> # x86_64, aarch64
Tested-by: Xi Ruoyao <xry111@xry111.site> # x86_64, aarch64, loongarch64
Tested-by: Stefan Liebler <stli@linux.ibm.com> # s390x
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Per the rseq syscall documentation, 3 fields are required to be
initialized by userspace prior to registration, they are 'cpu_id',
'rseq_cs' and 'flags'. Since we have no guarantee that 'struct pthread'
is cleared on all architectures, explicitly set those 3 fields prior to
registration.
Signed-off-by: Michael Jeanson <mjeanson@efficios.com>
Reviewed-by: Florian Weimer <fweimer@redhat.com>
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There are various existing tests that call pthread_attr_init and then
verify properties of the resulting initial values retrieved with
pthread_attr_get* functions. However, those are missing coverage of
the initial values retrieved with pthread_attr_getschedparam and
pthread_attr_getstacksize. Add testing for initial values from those
functions as well.
(tst-attr2 covers pthread_attr_getdetachstate,
pthread_attr_getguardsize, pthread_attr_getinheritsched,
pthread_attr_getschedpolicy, pthread_attr_getscope. tst-attr3 covers
some of those together with pthread_attr_getaffinity_np.
tst-pthread-attr-sigmask covers pthread_attr_getsigmask_np.
pthread_attr_getstack has unspecified results if called before the
relevant attributes have been set, while pthread_attr_getstackaddr is
deprecated.)
Tested for x86_64.
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The pthread_timedjoin_np and pthread_clockjoin_np functions do not
check that a valid time has been specified. The documentation for
these functions in the glibc manual isn't sufficiently detailed to say
if they should, but consistency with POSIX functions such as
pthread_mutex_timedlock and pthread_cond_timedwait strongly indicates
that an EINVAL error is appropriate (even if there might be some
ambiguity about exactly where such a check should go in relation to
other checks for whether the thread exists, whether it's immediately
joinable, etc.). Copy the logic for such a check used in
pthread_rwlock_common.c.
pthread_join_common had some logic calling valid_nanoseconds before
commit 9e92278ffad441daf588ff1ff5bd8094aa33fbfd, "nptl: Remove
clockwait_tid"; I haven't checked exactly what cases that detected.
Tested for x86_64 and x86.
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The recursive lock used on abort does not synchronize with a new process
creation (either by fork-like interfaces or posix_spawn ones), nor it
is reinitialized after fork().
Also, the SIGABRT unblock before raise() shows another race condition,
where a fork or posix_spawn() call by another thread, just after the
recursive lock release and before the SIGABRT signal, might create
programs with a non-expected signal mask. With the default option
(without POSIX_SPAWN_SETSIGDEF), the process can see SIG_DFL for
SIGABRT, where it should be SIG_IGN.
To fix the AS-safe, raise() does not change the process signal mask,
and an AS-safe lock is used if a SIGABRT is installed or the process
is blocked or ignored. With the signal mask change removal,
there is no need to use a recursive loc. The lock is also taken on
both _Fork() and posix_spawn(), to avoid the spawn process to see the
abort handler as SIG_DFL.
A read-write lock is used to avoid serialize _Fork and posix_spawn
execution. Both sigaction (SIGABRT) and abort() requires to lock
as writer (since both change the disposition).
The fallback is also simplified: there is no need to use a loop of
ABORT_INSTRUCTION after _exit() (if the syscall does not terminate the
process, the system is broken).
The proposed fix changes how setjmp works on a SIGABRT handler, where
glibc does not save the signal mask. So usage like the below will now
always abort.
static volatile int chk_fail_ok;
static jmp_buf chk_fail_buf;
static void
handler (int sig)
{
if (chk_fail_ok)
{
chk_fail_ok = 0;
longjmp (chk_fail_buf, 1);
}
else
_exit (127);
}
[...]
signal (SIGABRT, handler);
[....]
chk_fail_ok = 1;
if (! setjmp (chk_fail_buf))
{
// Something that can calls abort, like a failed fortify function.
chk_fail_ok = 0;
printf ("FAIL\n");
}
Such cases will need to use sigsetjmp instead.
The _dl_start_profile calls sigaction through _profil, and to avoid
pulling abort() on loader the call is replaced with __libc_sigaction.
Checked on x86_64-linux-gnu and aarch64-linux-gnu.
Reviewed-by: DJ Delorie <dj@redhat.com>
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Use the setresuid32 system call if it is available, prefering
it over setresuid. If both system calls exist, setresuid
is the 16-bit variant. This fixes a build failure on
sparcv9-linux-gnu.
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The current racy approach is to enable asynchronous cancellation
before making the syscall and restore the previous cancellation
type once the syscall returns, and check if cancellation has happen
during the cancellation entrypoint.
As described in BZ#12683, this approach shows 2 problems:
1. Cancellation can act after the syscall has returned from the
kernel, but before userspace saves the return value. It might
result in a resource leak if the syscall allocated a resource or a
side effect (partial read/write), and there is no way to program
handle it with cancellation handlers.
2. If a signal is handled while the thread is blocked at a cancellable
syscall, the entire signal handler runs with asynchronous
cancellation enabled. This can lead to issues if the signal
handler call functions which are async-signal-safe but not
async-cancel-safe.
For the cancellation to work correctly, there are 5 points at which the
cancellation signal could arrive:
[ ... )[ ... )[ syscall ]( ...
1 2 3 4 5
1. Before initial testcancel, e.g. [*... testcancel)
2. Between testcancel and syscall start, e.g. [testcancel...syscall start)
3. While syscall is blocked and no side effects have yet taken
place, e.g. [ syscall ]
4. Same as 3 but with side-effects having occurred (e.g. a partial
read or write).
5. After syscall end e.g. (syscall end...*]
And libc wants to act on cancellation in cases 1, 2, and 3 but not
in cases 4 or 5. For the 4 and 5 cases, the cancellation will eventually
happen in the next cancellable entrypoint without any further external
event.
The proposed solution for each case is:
1. Do a conditional branch based on whether the thread has received
a cancellation request;
2. It can be caught by the signal handler determining that the saved
program counter (from the ucontext_t) is in some address range
beginning just before the "testcancel" and ending with the
syscall instruction.
3. SIGCANCEL can be caught by the signal handler and determine that
the saved program counter (from the ucontext_t) is in the address
range beginning just before "testcancel" and ending with the first
uninterruptable (via a signal) syscall instruction that enters the
kernel.
4. In this case, except for certain syscalls that ALWAYS fail with
EINTR even for non-interrupting signals, the kernel will reset
the program counter to point at the syscall instruction during
signal handling, so that the syscall is restarted when the signal
handler returns. So, from the signal handler's standpoint, this
looks the same as case 2, and thus it's taken care of.
5. For syscalls with side-effects, the kernel cannot restart the
syscall; when it's interrupted by a signal, the kernel must cause
the syscall to return with whatever partial result is obtained
(e.g. partial read or write).
6. The saved program counter points just after the syscall
instruction, so the signal handler won't act on cancellation.
This is similar to 4. since the program counter is past the syscall
instruction.
So The proposed fixes are:
1. Remove the enable_asynccancel/disable_asynccancel function usage in
cancellable syscall definition and instead make them call a common
symbol that will check if cancellation is enabled (__syscall_cancel
at nptl/cancellation.c), call the arch-specific cancellable
entry-point (__syscall_cancel_arch), and cancel the thread when
required.
2. Provide an arch-specific generic system call wrapper function
that contains global markers. These markers will be used in
SIGCANCEL signal handler to check if the interruption has been
called in a valid syscall and if the syscalls has side-effects.
A reference implementation sysdeps/unix/sysv/linux/syscall_cancel.c
is provided. However, the markers may not be set on correct
expected places depending on how INTERNAL_SYSCALL_NCS is
implemented by the architecture. It is expected that all
architectures add an arch-specific implementation.
3. Rewrite SIGCANCEL asynchronous handler to check for both canceling
type and if current IP from signal handler falls between the global
markers and act accordingly.
4. Adjust libc code to replace LIBC_CANCEL_ASYNC/LIBC_CANCEL_RESET to
use the appropriate cancelable syscalls.
5. Adjust 'lowlevellock-futex.h' arch-specific implementations to
provide cancelable futex calls.
Some architectures require specific support on syscall handling:
* On i386 the syscall cancel bridge needs to use the old int80
instruction because the optimized vDSO symbol the resulting PC value
for an interrupted syscall points to an address outside the expected
markers in __syscall_cancel_arch. It has been discussed in LKML [1]
on how kernel could help userland to accomplish it, but afaik
discussion has stalled.
Also, sysenter should not be used directly by libc since its calling
convention is set by the kernel depending of the underlying x86 chip
(check kernel commit 30bfa7b3488bfb1bb75c9f50a5fcac1832970c60).
* mips o32 is the only kABI that requires 7 argument syscall, and to
avoid add a requirement on all architectures to support it, mips
support is added with extra internal defines.
Checked on aarch64-linux-gnu, arm-linux-gnueabihf, powerpc-linux-gnu,
powerpc64-linux-gnu, powerpc64le-linux-gnu, i686-linux-gnu, and
x86_64-linux-gnu.
[1] https://lkml.org/lkml/2016/3/8/1105
Reviewed-by: Carlos O'Donell <carlos@redhat.com>
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Fix an issue with commit b74121ae4bc5 ("Update.") and prevent a stray
process from being left behind by tst-cancel7 (and also tst-cancelx7,
which is the same test built with '-fexceptions' additionally supplied
to the compiler), which then blocks remote testing until the process has
been killed by hand.
This test case creates a thread that runs an extra copy of the test via
system(3) and using the '--direct' option so that the test wrapper does
not interfere with this instance. This extra copy executes its business
and calls sigsuspend(2) and then never terminates by itself. Instead it
relies on being killed by the main test process directly via a thread
cancellation request or, should that fail, by issuing SIGKILL either at
the conclusion of 'do_test' or by the test driver via 'do_cleanup' where
the test timeout has been hit or the test driver interrupted.
However if the main test process has been instead killed by a signal,
such as due to incorrect execution, before it had a chance to kill the
extra copy of the test case, then the test wrapper will terminate
without running 'do_cleanup' and consequently the extra copy of the test
case will remain forever in its suspended state, and in the remote case
in particular it means that the remote test wrapper will wait forever
for the SSH command to complete.
This has been observed with the 'alpha-linux-gnu' target, where the main
test process triggers SIGSEGV and the test wrapper correctly records:
Didn't expect signal from child: got `Segmentation fault'
in nptl/tst-cancel7.out and terminates, but then the calling SSH command
continues waiting for the remaining process started in the same session
on the remote target to complete.
Address this problem by also registering 'do_cleanup' via atexit(3),
observing that 'support_delete_temp_files' is registered by the test
wrapper before the test initializing function 'do_prepare' is called and
that we call all the functions registered in the reverse of the order in
which they were registered, so it is safe to refer to 'pidfilename' in
'do_cleanup' invoked by exit(3) because by that time temporary files
have not yet been deleted.
A minor inconvenience is that if 'signal_handler' is invoked in the test
wrapper as a result of SIGALRM rather than SIGINT, then 'do_cleanup'
will be called twice, once as a cleanup handler and again by exit(3).
In reality it is harmless though, because issuing SIGKILL is guarded by
a record lock, so if the first call has succeeded in killing the extra
copy of the test case, then the subsequent call will do nothing.
Reviewed-by: Adhemerval Zanella <adhemerval.zanella@linaro.org>
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Move the release of the semaphore used to synchronize between an extra
copy of the test run as a separate process and the main test process
until after the PID file has been locked. It is so that if the cleanup
function gets called by the test driver due to premature termination of
the main test process, then the function does not get at the PID file
before it has been locked and conclude that the extra copy of the test
has already terminated. This won't usually happen due to a relatively
high amount of time required to elapse before timeout triggers in the
test driver, but it will change with the next change.
There is still a small time window remaining with this change in place
where the main test process gets killed for some reason between the
extra copy of the test has been already started by pthread_create(3) and
a successful return from the call to sem_wait(3), in which case the
cleanup function can be reached before PID has been written to the PID
file and the file locked. It seems that with the test case structured
as it is now and PID-based process management we have no means to avoid
it.
Reviewed-by: Adhemerval Zanella <adhemerval.zanella@linaro.org>
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Also remove an outdated comment: _dl_allocate_tls_init is
called as part of pthread_create.
Reviewed-by: Carlos O'Donell <carlos@redhat.com>
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Fix an issue with commit 2af4e3e5668f ("Test of semaphores.") by making
the tst-sem11 and tst-sem12 tests use the test driver, preventing them
from ever causing testing to hang forever and never complete, such as
currently happening with the 'mips-linux-gnu' (o32 ABI) target. Adjust
the name of the PREPARE macro, which clashes with the interpretation of
its presence by the test driver, by using a TF_ prefix in reference to
the name of the 'tf' function.
Reviewed-by: Adhemerval Zanella <adhemerval.zanella@linaro.org>
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Add a copyright notice to the tst-sem11 and tst-sem12 tests, observing
that they have been originally contributed back in 2007, with commit
2af4e3e5668f ("Test of semaphores.").
Reviewed-by: Adhemerval Zanella <adhemerval.zanella@linaro.org>
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