this is analogous commit fffc5cda
which fixed the corresponding issue for mutexes.

the robust list can't be used here because the locks do not share a
common layout with mutexes. at some point it may make sense to simply
incorporate a mutex object into the FILE structure and use it, but
that would be a much more invasive change, and it doesn't mesh well
with the current design that uses a simpler code path for internal
locking and pulls in the recursive-mutex-like code when the flockfile
API is used explicitly.

for unknown syscall commands, the kernel produces ENOSYS, not EINVAL.

the subsequent code in pthread_create and the code which copies TLS
initialization images to the new thread's TLS space assume that the
memory provided to them is zero-initialized, which is true when it's
obtained by pthread_create using mmap. however, when the caller
provides a stack using pthread_attr_setstack, pthread_create cannot
make any assumptions about the contents. simply zero-filling the
relevant memory in this case is the simplest and safest fix.

unfortunately this needs to be able to vary by arch, because of a huge
mess GCC made: the GCC definition, which became the ABI, depends on
quirks in GCC's definition of __alignof__, which does not match the
formal alignment of the type.

GCC's __alignof__ unexpectedly exposes the an implementation detail,
its "preferred alignment" for the type, rather than the formal/ABI
alignment of the type, which it only actually uses in structures. on
most archs the two values are the same, but on some (at least i386)
the preferred alignment is greater than the ABI alignment.

I considered using _Alignas(8) unconditionally, but on at least one
arch (or1k), the alignment of max_align_t with GCC's definition is
only 4 (even the "preferred alignment" for these types is only 4).

the main idea of the changes made is to have waiters wait directly on
the "barrier" lock that was used to prevent them from making forward
progress too early rather than first waiting on the atomic state value
and then attempting to lock the barrier.

in addition, adjustments to the mutex waiter count are optimized.
previously, each waking waiter decremented the count (unless it was
the first) then immediately incremented it again for the next waiter
(unless it was the last). this was a roundabout was of achieving the
equivalent of incrementing it once for the first waiter and
decrementing it once for the last.

previously, wake order could be unpredictable: if a waiter happened to
leave its futex wait on the state early, e.g. due to EAGAIN while
restarting after a signal handler, it could acquire the mutex out of
turn. handling this required ugly O(n) list walking in the unwait
function and accounting to remove waiters that already woke from the
list.

with the new changes, the "barrier" locks in each waiter node are only
unlocked in turn. in addition to simplifying the code, this seems to
improve performance slightly, probably by reducing the number of
accesses threads make to each other's stacks.

as an additional benefit, unrecoverable mutex re-locking errors
(mainly ENOTRECOVERABLE for robust mutexes) no longer need to be
handled with deadlock; they can be reported to the caller, since the
unlocking sequence makes it unnecessary to rely on the mutex to
synchronize access to the waiter list.

the immediate issue that was reported by Jens Gustedt and needed to be
fixed was corruption of the cv/mutex waiter states when switching to
using a new mutex with the cv after all waiters were unblocked but
before they finished returning from the wait function.

self-synchronized destruction was also handled poorly and may have had
race conditions. and the use of sequence numbers for waking waiters
admitted a theoretical missed-wakeup if the sequence number wrapped
through the full 32-bit space.

the new implementation is largely documented in the comments in the
source. the basic principle is to use linked lists initially attached
to the cv object, but detachable on signal/broadcast, made up of nodes
residing in automatic storage (stack) on the threads that are waiting.
this eliminates the need for waiters to access the cv object after
they are signaled, and allows us to limit wakeup to one waiter at a
time during broadcasts even when futex requeue cannot be used.

performance is also greatly improved, roughly double some tests.

basically nothing is changed in the process-shared cond var case,
where this implementation does not work, since processes do not have
access to one another's local storage.

when the kernel is responsible for waking waiters on a robust mutex
whose owner died, it does not have a waiters count available and must
rely entirely on the waiter bit of the lock value.

normally, this bit is only set by newly arriving waiters, so it will
be clear if no new waiters arrived after the current owner obtained
the lock, even if there are other waiters present. leaving it clear is
desirable because it allows timed-lock operations to remove themselves
as waiters and avoid causing unnecessary futex wake syscalls. however,
for process-shared robust mutexes, we need to set the bit whenever
there are existing waiters so that the kernel will know to wake them.

for non-process-shared robust mutexes, the wake happens in userspace
and can look at the waiters count, so the bit does not need to be set
in the non-process-shared case.

when manipulating the robust list, the order of stores matters,
because the code may be asynchronously interrupted by a fatal signal
and the kernel will then access the robust list in what is essentially
an async-signal context.

previously, aliasing considerations made it seem unlikely that a
compiler could reorder the stores, but proving that they could not be
reordered incorrectly would have been extremely difficult. instead
I've opted to make all the pointers used as part of the robust list,
including those in the robust list head and in the individual mutexes,
volatile.

in addition, the format of the robust list has been changed to point
back to the head at the end, rather than ending with a null pointer.
this is to match the documented kernel robust list ABI. the null
pointer, which was previously used, only worked because faults during
access terminate the robust list processing.

a robust mutex should not enter the unrecoverable status until it's
unlocked without marking it consistent. previously, flag 8 in the type
was used as an indication of unrecoverable, but only honored after
successful locking; this resulted in a race window where the
unrecoverable mutex could appear to a second thread as locked/busy
again while the first thread was in the process of observing it as
unrecoverable.

now, flag 8 is used to mean that the mutex is in the process of being
recovered, but not yet marked consistent. the flag only takes effect
in pthread_mutex_unlock, where it causes the value 0x40000000 (owner
dead flag, with old owner tid 0, an otherwise impossible state) to be
stored in the lock. subsequent lock attempts will interpret this state
as unrecoverable.

per the resolution of Austin Group issue 755, the POSIX requirement
that ownership be enforced for recursive and error-checking mutexes
does not allow a random new thread to acquire ownership of an orphaned
mutex just because it happened to be assigned the same tid as the
original owner that exited with the mutex locked.

one possible fix for this issue would be to disallow the kernel thread
to terminate when it exited with mutexes held, permanently reserving
the tid against reuse. however, this does not solve the problem for
process-shared mutexes where lifetime cannot be controlled, so it was
not used.

the alternate approach I've taken is to reuse the robust mutex system
for non-robust recursive and error-checking mutexes. when a thread
exits, the kernel (or the new userspace robust-list code added in
commit b092f1c5) will set the
owner-died bit for these orphaned mutexes, but since the mutex-type is
not robust, pthread_mutex_trylock will not allow a new owner to
acquire them. instead, they remain in a state of being permanently
locked, as desired.

the whole point of this locking is to prevent munmap, or mmap with
MAP_FIXED, from deallocating virtual addresses, or changing the
backing a given virtual address refers to, during certain race windows
involving self-synchronized unmapping or destruction of pthread
synchronization objects. there is no need for exclusion in the other
direction, so it suffices to take the lock momentarily and release it
before making the syscall, rather than holding it across the syscall.

the kernel always uses non-private wake when walking the robust list
when a thread or process exits, so it's not able to wake waiters
listening with the private futex flag. this problem is solved by doing
the equivalent in userspace as the last step of pthread_exit.

care is taken to remove mutexes from the robust list before unlocking
them so that the kernel will not attempt to access them again,
possibly after another thread locks them. this removal code can treat
the list as singly-linked, since no further code which would add or
remove items is able to run at this point. moreover, the pending
pointer is not needed since the mutexes being unlocked are all
process-local; in the case of asynchronous process termination, they
all cease to exist.

since a process-local robust mutex cannot come into existence without
a call to pthread_mutexattr_setrobust in the same process, the code
for userspace robust list processing is put in that source file, and
a weak alias to a dummy function is used to avoid pulling in this
bloat as part of pthread_exit in static-linked programs.

private-futex uses the virtual address of the futex int directly as
the hash key rather than requiring the kernel to resolve the address
to an underlying backing for the mapping in which it lies. for certain
usage patterns it improves performance significantly.

in many places, the code using futex __wake and __wait operations was
already passing a correct fixed zero or nonzero flag for the priv
argument, so no change was needed at the site of the call, only in the
__wake and __wait functions themselves. in other places, especially
where the process-shared attribute for a synchronization object was
not previously tracked, additional new code is needed. for mutexes,
the only place to store the flag is in the type field, so additional
bit masking logic is needed for accessing the type.

for non-process-shared condition variable broadcasts, the futex
requeue operation is unable to requeue from a private futex to a
process-shared one in the mutex structure, so requeue is simply
disabled in this case by waking all waiters.

for robust mutexes, the kernel always performs a non-private wake when
the owner dies. in order not to introduce a behavioral regression in
non-process-shared robust mutexes (when the owning thread dies), they
are simply forced to be treated as process-shared for now, giving
correct behavior at the expense of performance. this can be fixed by
adding explicit code to pthread_exit to do the right thing for
non-shared robust mutexes in userspace rather than relying on the
kernel to do it, and will be fixed in this way later.

since not all supported kernels have private futex support, the new
code detects EINVAL from the futex syscall and falls back to making
the call without the private flag. no attempt to cache the result is
made; caching it and using the cached value efficiently is somewhat
difficult, and not worth the complexity when the benefits would be
seen only on ancient kernels which have numerous other limitations and
bugs anyway.

C99 6.10.3p11 disallows such constructs
so use an #ifdef outside of the argument list of __syscall

isspace can be a bottleneck in a simple parser, inlining it
gives slightly smaller and faster code

src/locale/pleval.o already had this optimization, the size
change for other libc functions for i386 is

src/internal/intscan.o     2134    2118   -16
src/locale/dcngettext.o    1562    1552   -10
src/network/res_msend.o    1961    1940   -21
src/network/lookup_name.o  2627    2608   -19
src/network/getnameinfo.o  1814    1811    -3
src/network/lookup_serv.o   643     624   -19
src/stdio/vfscanf.o        2675    2663   -12
src/stdlib/atoll.o          117     107   -10
src/stdlib/atoi.o            95      91    -4
src/stdlib/atol.o            95      91    -4
src/time/strptime.o        1515    1503   -12
(TOTALS)                 432451  432321  -130

when the dynamic loader is disabled, dlopen fails correctly but dlerror
did not return a human readable error string like it should have.

glibc declares clearenv under _BSD_SOURCE, some applications might
depend on it being available this way.

The function originates from SunOS 4.x in which the null argument
is allowed. glibc also handles this case.

the code which loads locale files was already rejecting locale names
containing slashes. however, LC_MESSAGES records a locale name even if
libc does not have a matching locale file, so that gettext or
application code can use the recorded locale name for message
translations to languages that libc does not support. this recorded
name was not being checked for slashes, meaning that such code could
potentially be tricked into directory traversal.

in addition, since the value of a locale category is sometimes used as
a pathname component by callers, the improved code rejects any value
beginning with a dot. this prevents traversal to the parent directory
via "..", use of the top-level locale directory via ".", and also
avoids "hidden" directories as a side effect.

finally, overly long locale names are now rejected (treated as an
unrecognized name and thus as an alias for C.UTF-8) rather than being
truncated.

per the resolution of Austin Group issue #617, these are accepted for
XSI option in POSIX future and thus I'm treating them as standard
functions.

per the standard, ffs is XSI shaded, whereas the other functions in
this header are in the base.

for or1k, the kernel expects the offset passed to mmap2 in units of
the 8k page size, not the standard unit of 4k used on most other
archs.

according to Stefan Kristiansson, or1k page size is not actually
variable and the value of 8192 is part of the ABI.

using an operator precedence parser the code size
became smaller and it is only slower by about %10

size of old vs new pleval.o on different archs:
(with inlined isspace added to pleval.c for now)

old:
   text    data     bss     dec     hex filename
    828       0       0     828     33c pl.i386.o
   1152       0       0    1152     480 pl.arm.o
   1704       0       0    1704     6a8 pl.mips.o
   1328       0       0    1328     530 pl.ppc.o
    992       0       0     992     3e0 pl.x64.o
new:
   text    data     bss     dec     hex filename
    693       0       0     693     2b5 pl.i386.o
    972       0       0     972     3cc pl.arm.o
   1276       0       0    1276     4fc pl.mips.o
   1087       0       0    1087     43f pl.ppc.o
    846       0       0     846     34e pl.x64.o

the previous implementations had several deficiencies, the most severe
of which was the inability to report unconfigured interfaces or
interfaces without ipv4 addresses. among the options discussed for
fixing this, using netlink turned out to be the one with the least
cost and most additional advantages. other improvements include:

if_nameindex now avoids duplicates in the list it produces, but still
includes legacy-style interface aliases if any are in use.

getifaddrs now reports hardware addresses and includes the scope_id
for link-local ipv6 addresses in the resulting address.

this commit changes the names to match the kernel names, exposing
under the normal names the "old" versions which work with a smaller
termios structure compatible with the userspace structure, and
renaming the "new" versions with "2" on the end like the kernel has.

this fixes spurious warnings "Unsupported ioctl: cmd=0x802c542a" from
qemu-sh4 and should be more correct anyway, since our userspace
termios structure does not have meaningful information in the part
which the kernel would be interpreting as speeds with the new ioctl.

const parsing, depth accounting and failure handling was changed
a bit so the generated code is slightly smaller.

while the __mo_lookup backend can verify that the translated message
ends with a null terminator, is has no way to know nplurals and thus
no way to verify that sufficiently many null terminators are present
in the string to satisfy all plural forms. the code in dcngettext was
already attempting to avoid reading past the end of the mo file
mapping, but failed to do so because the strlen call itself could
over-read. using strnlen instead allows us to avoid the problem.

rather than just checking that the start of the string lies within the
mapping, also check that the nominal length remains within the
mapping, and that the null terminator is present at the nominal
length. this ensures that the caller, using the result as a C string,
will not read past the end of the mapping.

the nominal length is never exposed to the caller, but it's useful
internally to find where the null terminator should be without having
to restort to linear search via strnlen/memchr.

the new code in dcngettext was written by me, and the expression
evaluator by Szabolcs Nagy (nsz).

this follows the same logic as in the previous commit for other archs.

the a_cas_l, a_swap_l, a_swap_p, and a_store_l operations were
probably used a long time ago when only i386 and x86_64 were
supported. as other archs were added, support for them was
inconsistent, and they are obviously not in use at present. having
them around potentially confuses readers working on new ports, and the
type-punning hacks and inconsistent use of types in their definitions
is not a style I wish to perpetuate in the source tree, so removing
them seems appropriate.

while other usage I've seen only has the synco instruction after the
atomic operation, I cannot find any documentation indicating that this
is correct. certainly all stores before the atomic need to have been
synchronized before the atomic operation takes place.