linux's sched_* syscalls actually implement the TPS (thread
scheduling) functionality, not the PS (process scheduling)
functionality which the sched_* functions are supposed to have.
omitting support for the PS option (and having the sched_* interfaces
fail with ENOSYS rather than omitting them, since some broken software
assumes they exist) seems to be the only conforming way to do this on
linux.
this mirrors the stdio_impl.h cleanup. one header which is not
strictly needed, errno.h, is left in pthread_impl.h, because since
pthread functions return their error codes rather than using errno,
nearly every single pthread function needs the errno constants.
in a few places, rather than bringing in string.h to use memset, the
memset was replaced by direct assignment. this seems to generate much
better code anyway, and makes many functions which were previously
non-leaf functions into leaf functions (possibly eliminating a great
deal of bloat on some platforms where non-leaf functions require ugly
prologue and/or epilogue).
despite documentation that makes it sound a lot different, the only
ABI-constraint difference between TLS variants II and I seems to be
that variant II stores the initial TLS segment immediately below the
thread pointer (i.e. the thread pointer points to the end of it) and
variant I stores the initial TLS segment above the thread pointer,
requiring the thread descriptor to be stored below. the actual value
stored in the thread pointer register also tends to have per-arch
random offsets applied to it for silly micro-optimization purposes.
with these changes applied, TLS should be basically working on all
supported archs except microblaze. I'm still working on getting the
necessary information and a working toolchain that can build TLS
binaries for microblaze, but in theory, static-linked programs with
TLS and dynamic-linked programs where only the main executable uses
TLS should already work on microblaze.
alignment constraints have not yet been heavily tested, so it's
possible that this code does not always align TLS segments correctly
on archs that need TLS variant I.
the code in __libc_start_main is now responsible for parsing auxv,
rather than duplicating the parsing all over the place. this should
shave off a few cycles and some code size. __init_libc is left as an
external-linkage function despite the fact that it could be static, to
prevent it from being inlined and permanently wasting stack space when
main is called.
a few other minor changes are included, like eliminating per-thread
ssp canaries (they were likely broken when combined with certain
dlopen usages, and completely unnecessary) and some other unnecessary
checks. since this code gets linked into every program, it should be
as small and simple as possible.
unlike other implementations, this one reserves memory for new TLS in
all pre-existing threads at dlopen-time, and dlopen will fail with no
resources consumed and no new libraries loaded if memory is not
available. memory is not immediately distributed to running threads;
that would be too complex and too costly. instead, assurances are made
that threads needing the new TLS can obtain it in an async-signal-safe
way from a buffer belonging to the dynamic linker/new module (via
atomic fetch-and-add based allocator).
I've re-appropriated the lock that was previously used for __synccall
(synchronizing set*id() syscalls between threads) as a general
pthread_create lock. it's a "backwards" rwlock where the "read"
operation is safe atomic modification of the live thread count, which
multiple threads can perform at the same time, and the "write"
operation is making sure the count does not increase during an
operation that depends on it remaining bounded (__synccall or dlopen).
in static-linked programs that don't use __synccall, this lock is a
no-op and has no cost.
this code will not work yet because the necessary relocations are not
supported, and cannot be supported without some internal changes to
how relocation processing works (coming soon).
the design for TLS in dynamic-linked programs is mostly complete too,
but I have not yet implemented it. cost is nonzero but still low for
programs which do not use TLS and/or do not use threads (a few hundred
bytes of new code, plus dependency on memcpy). i believe it can be
made smaller at some point by merging __init_tls and __init_security
into __libc_start_main and avoiding duplicate auxv-parsing code.
at the same time, I've also slightly changed the logic pthread_create
uses to allocate guard pages to ensure that guard pages are not
counted towards commit charge.
this function was overly complicated and not even obviously correct.
avoid using openat/linkat just like in shm_open, and instead expand
pathname using code shared with shm_open. remove bogus (and dangerous,
with priorities) use of spinlocks.
this commit also heavily streamlines the code and ensures there are no
failure cases that can happen after a new semaphore has been created
in the filesystem, since that case is unreportable.
based on initial work by rdp, with heavy modifications. some features
including threads are untested because qemu app-level emulation seems
to be broken and I do not have a proper system image for testing.
note that POSIX does not specify these functions as _Noreturn, because
POSIX is aligned with C99, not the new C11 standard. when POSIX is
eventually updated to C11, it will almost surely give these functions
the _Noreturn attribute. for now, the actual _Noreturn keyword is not
used anyway when compiling with a c99 compiler, which is what POSIX
requires; the GCC __attribute__ is used instead if it's available,
however.
in a few places, I've added infinite for loops at the end of _Noreturn
functions to silence compiler warnings. presumably
__buildin_unreachable could achieve the same thing, but it would only
work on newer GCCs and would not be portable. the loops should have
near-zero code size cost anyway.
like the previous _Noreturn commit, this one is based on patches
contributed by philomath.
to deal with the fact that the public headers may be used with pre-c99
compilers, __restrict is used in place of restrict, and defined
appropriately for any supported compiler. we also avoid the form
[restrict] since older versions of gcc rejected it due to a bug in the
original c99 standard, and instead use the form *restrict.
if new shared mappings of files/devices/shared memory can be made
between the time a robust mutex is unlocked and its subsequent removal
from the pending slot in the robustlist header, the kernel can
inadvertently corrupt data in the newly-mapped pages when the process
terminates. i am fixing the bug by using the same global vm lock
mechanism that was used to fix the race condition with unmapping
barriers after pthread_barrier_wait returns.
some minor changes to how hard-coded sets for thread-related purposes
are handled were also needed, since the old object sizes were not
necessarily sufficient. things have gotten a bit ugly in this area,
and i think a cleanup is in order at some point, but for now the goal
is just to get the code working on all supported archs including mips,
which was badly broken by linux rejecting syscalls with the wrong
sigset_t size.
this seems counter-intuitive since sem_trywait is supposed to just try
once, not wait for the semaphore. however, the retry loop is not a
wait. instead, it's to handle the case where the value changes due to
a simultaneous post or wait from another thread while the semaphore
value remains positive. in such a case, it's absolutely wrong for
sem_trywait to fail with EAGAIN because the semaphore is not busy.
these could have caused memory corruption due to invalid accesses to
the next field. all should be fixed now; I found the errors with fgrep
-r '__lock(&', which is bogus since the argument should be an array.
after the thread unmaps its own stack/thread structure, the kernel,
performing child tid clear and futex wake, could clobber a new mapping
made at the same location as the just-removed thread's tid field.
disable kernel clearing of child tid to prevent this.
the mips abi reserves stack space equal to the size of the in-register
args for the callee to save the args, if desired. this would cause the
beginning of the thread structure to be clobbered...
the old code worked in qemu app-level emulation, but not on real
kernels where the clone syscall does not copy the register values to
the new thread. save arguments on the new thread stack instead.
i originally omitted these (optional, per POSIX) interfaces because i
considered them backwards implementation details. however, someone
later brought to my attention a fairly legitimate use case: allocating
thread stacks in memory that's setup for sharing and/or fast transfer
between CPU and GPU so that the thread can move data to a GPU directly
from automatic-storage buffers without having to go through additional
buffer copies.
perhaps there are other situations in which these interfaces are
useful too.
this fix is necessary because a program could be started with some of
the implementation-reserved signals masked (e.g. due to exec having
been called from a signal handler, or from a non-musl program) and
then could obtain an invalid-to-use-later sigset_t as the old/saved
signal mask.
this action is now performed in pthread_self initialization; it must
be performed there in case the first call to pthread_create is from a
signal handler, in which case the old signal mask could be restored on
return from the signal.
if the process started with these signals blocked, cancellation could
fail or setxid could deadlock. there is no way to globally unblock
them after threads have been created. by unblocking them in the
pthread_self initialization for the main thread, we ensure that
they're unblocked before any other threads are created and also
outside of any signal handler context (sigaction initialized
pthread_self), which is important so that return from a signal handler
won't re-block them.
the bug was that cancellation requests which arrived while a
cancellation point was interrupted by a signal handler would not be
acted upon when the signal handler returns. this was because cp_sp was
never set; it's no longer needed or used.
instead, just always re-raise the signal when cancellation was not
acted upon. this wastes a tiny amount of time in the rare case where
it even matters, but it ensures correctness and simplifies the code.
stale state information indicating that a thread was possibly blocked
at a cancellation point could get left behind if longjmp was used to
exit a signal handler that interrupted a cancellation point.
to fix the issue, we throw away the state information entirely and
simply compare the saved instruction pointer to a range of code
addresses in the __syscall_cp_asm function. all the ugly PIC work
(which becomes minimal anyway with this approach) is defered to
cancellation time instead of happening at every syscall, which should
improve performance too.
this commit also fixes cancellation on arm, which was mildly broken
(race condition, not checking cancellation flag once inside the
cancellation point zone). apparently i forgot to implement that. the
new arm code is untested, but appears correct; i'll test and fix it
later if there are problems.
i made a best attempt, but the intended semantics of this function are
fundamentally contradictory. there is no consistent way to handle
ownership of locks when forking a multi-threaded process. the code
could have worked by accident for programs that only used normal
mutexes and nothing else (since they don't actually store or care
about their owner), but that's about it. broken-by-design interfaces
that aren't even in glibc (only solaris) don't belong in musl.
these changes are based on the following communication via email:
"I hereby grant that all of the code I have contributed to musl on or
before April 23, 2012 may be licensed under the terms of the following
MIT license:
Copyright (c) 2011-2012 Nicholas J. Kain
Permission is hereby granted, free of charge, to any person obtaining
a copy of this software and associated documentation files (the
"Software"), to deal in the Software without restriction, including
without limitation the rights to use, copy, modify, merge, publish,
distribute, sublicense, and/or sell copies of the Software, and to
permit persons to whom the Software is furnished to do so, subject to
the following conditions:
The above copyright notice and this permission notice shall be
included in all copies or substantial portions of the Software.
THE SOFTWARE IS PROVIDED "AS IS", WITHOUT WARRANTY OF ANY KIND,
EXPRESS OR IMPLIED, INCLUDING BUT NOT LIMITED TO THE WARRANTIES OF
MERCHANTABILITY, FITNESS FOR A PARTICULAR PURPOSE AND NONINFRINGEMENT.
IN NO EVENT SHALL THE AUTHORS OR COPYRIGHT HOLDERS BE LIABLE FOR ANY
CLAIM, DAMAGES OR OTHER LIABILITY, WHETHER IN AN ACTION OF CONTRACT,
TORT OR OTHERWISE, ARISING FROM, OUT OF OR IN CONNECTION WITH THE
SOFTWARE OR THE USE OR OTHER DEALINGS IN THE SOFTWARE."
this change is necessary or pthread_create will always fail on
security-hardened kernels. i considered first trying to make the stack
executable and simply retrying without execute permissions when the
first try fails, but (1) this would incur a serious performance
penalty on hardened systems, and (2) having the stack be executable is
just a bad idea from a security standpoint.
if there is real-world "GNU C" code that uses nested functions with
threads, and it can't be fixed, we'll have to consider other ways of
solving the problem, but for now this seems like the best fix.
pthread structure has been adjusted to match the glibc/GCC abi for
where the canary is stored on i386 and x86_64. it will need variants
for other archs to provide the added security of the canary's entropy,
but even without that it still works as well as the old "minimal" ssp
support. eventually such changes will be made anyway, since they are
also needed for GCC/C11 thread-local storage support (not yet
implemented).
care is taken not to attempt initializing the thread pointer unless
the program actually uses SSP (by reference to __stack_chk_fail).
i did some testing trying to switch malloc to use the new internal
lock with priority inheritance, and my malloc contention test got
20-100 times slower. if priority inheritance futexes are this slow,
it's simply too high a price to pay for avoiding priority inversion.
maybe we can consider them somewhere down the road once the kernel
folks get their act together on this (and perferably don't link it to
glibc's inefficient lock API)...
as such, i've switch __lock to use malloc's implementation of
lightweight locks, and updated all the users of the code to use an
array with a waiter count for their locks. this should give optimal
performance in the vast majority of cases, and it's simple.
malloc is still using its own internal copy of the lock code because
it seems to yield measurably better performance with -O3 when it's
inlined (20% or more difference in the contention stress test).
this bug probably would have gone unnoticed since it's only used in
the fallback code for systems where priority-inheritance locking
fails. unfortunately this approach results in one spurious wake
syscall on the final unlock, when there are no waiters remaining. the
alternative (possibly better) would be to use broadcast wakes instead
of reflagging the waiter unconditionally, and let each waiter reflag
itself; this saves one syscall at the expense of invoking the
"thundering herd" effect (worse performance degredation) when there
are many waiters.
ideally we would be able to update all of our locks to use an array of
two ints rather than a single int, and use a separate counter system
like proper mutexes use; then we could avoid all spurious wake calls
without resorting to broadcasts. however, it's not clear to me that
priority inheritance futexes support this usage. the kernel sets the
waiters flag for them (just like we're doing now) and i can't tell if
it's safe to bypass the kernel when unlocking just because we know
(from private data, the waiter count) that there are no waiters. this
is something that could be explored in the future.
we use priority inheritance futexes if possible so that the library
cannot hit internal priority inversion deadlocks in the presence of
realtime priority scheduling (full support to be added later).
Rich Felker