The overhead of unconditionally allocating TIDs (and likewise,
unconditionally deallocating them), is amortized across multiple
thread creations by the way UMA makes it possible to have type-stable
storage.
Previously the cost was kept down by having threads created as part
of a fork operation use the process' PID as the TID. While this had
some nice properties, it also introduced complexity in the way TIDs
were allocated. Most importantly, by using the type-stable storage
that UMA gives us this was also unnecessary.
This change affects how core dumps are created and in particular how
the PRSTATUS notes are dumped. Since we don't have a thread with a
TID equalling the PID, we now need a different way to preserve the
old and previous behavior. We do this by having the given thread (i.e.
the thread passed to the core dump code in td) dump it's state first
and fill in pr_pid with the actual PID. All other threads will have
pr_pid contain their TIDs. The upshot of all this is that the debugger
will now likely select the right LWP (=TID) as the initial thread.
Credits to: julian@ for spotting how we can utilize UMA.
Thanks to: all who provided julian@ with test results.
is twofold:
1. When a 1:1 or M:N threaded process dumps core, we need to put the
register state of each of its kernel threads in the core file.
This can only be done by differentiating the pid field in the
respective note. For this we need the tid.
2. When thread support is present for remote debugging the kernel
with gdb(1), threads need to be identified by an integer due to
limitations in the remote protocol. This requires having a tid.
To minimize the impact of having thread IDs, threads that are created
as part of a fork (i.e. the initial thread in a process) will inherit
the process ID (i.e. tid=pid). Subsequent threads will have IDs larger
than PID_MAX to avoid interference with the pid allocation algorithm.
The assignment of tids is handled by thread_new_tid().
The thread ID allocation algorithm has been written with 3 assumptions
in mind:
1. IDs need to be created as fast a possible,
2. Reuse of IDs may happen instantaneously,
3. Someone else will write a better algorithm.
in the two consumers that need it.. processes using AIO and netncp.
Update docs. Say that process_exec is called with Giant, but not to
depend on it. All our consumers can handle it without Giant.
- no longer serialize on Giant for thread_single*() and family in fork,
exit and exec
- thread_wait() is mpsafe, assert no Giant
- reduce scope of Giant in exit to not cover thread_wait and just do
vm_waitproc().
- assert that thread_single() family are not called with Giant
- remove the DROP/PICKUP_GIANT macros from thread_single() family
- assert that thread_suspend_check() s not called with Giant
- remove manual drop_giant hack in thread_suspend_check since we know it
isn't held.
- remove the DROP/PICKUP_GIANT macros from thread_suspend_check() family
- mark kse_create() mpsafe
that we (p1) are currently running, we hold a reference on p_textvp which
means the vnode cannot go away. p2 cannot run yet (and hence cannot exit)
so this should be safe to do at this point. As a bonus, it removes a
block of under-Giant code that was there to support the vref.
process group and session dereferences. Also, check that p_pgrp and
p_sesssion are NULL before dereferencing them.
- Push down Giant in fork1().
Requested by: peter
the syscall arguments and does the suser() permission check, and
kern_mlock(), which does the resource limit checking and calls
vm_map_wire(). Split munlock() in a similar way.
Enable the RLIMIT_MEMLOCK checking code in kern_mlock().
Replace calls to vslock() and vsunlock() in the sysctl code with
calls to kern_mlock() and kern_munlock() so that the sysctl code
will obey the wired memory limits.
Nuke the vslock() and vsunlock() implementations, which are no
longer used.
Add a member to struct sysctl_req to track the amount of memory
that is wired to handle the request.
Modify sysctl_wire_old_buffer() to return an error if its call to
kern_mlock() fails. Only wire the minimum of the length specified
in the sysctl request and the length specified in its argument list.
It is recommended that sysctl handlers that use sysctl_wire_old_buffer()
should specify reasonable estimates for the amount of data they
want to return so that only the minimum amount of memory is wired
no matter what length has been specified by the request.
Modify the callers of sysctl_wire_old_buffer() to look for the
error return.
Modify sysctl_old_user to obey the wired buffer length and clean up
its implementation.
Reviewed by: bms
- struct plimit includes a mutex to protect a reference count. The plimit
structure is treated similarly to struct ucred in that is is always copy
on write, so having a reference to a structure is sufficient to read from
it without needing a further lock.
- The proc lock protects the p_limit pointer and must be held while reading
limits from a process to keep the limit structure from changing out from
under you while reading from it.
- Various global limits that are ints are not protected by a lock since
int writes are atomic on all the archs we support and thus a lock
wouldn't buy us anything.
- All accesses to individual resource limits from a process are abstracted
behind a simple lim_rlimit(), lim_max(), and lim_cur() API that return
either an rlimit, or the current or max individual limit of the specified
resource from a process.
- dosetrlimit() was renamed to kern_setrlimit() to match existing style of
other similar syscall helper functions.
- The alpha OSF/1 compat layer no longer calls getrlimit() and setrlimit()
(it didn't used the stackgap when it should have) but uses lim_rlimit()
and kern_setrlimit() instead.
- The svr4 compat no longer uses the stackgap for resource limits calls,
but uses lim_rlimit() and kern_setrlimit() instead.
- The ibcs2 compat no longer uses the stackgap for resource limits. It
also no longer uses the stackgap for accessing sysctl's for the
ibcs2_sysconf() syscall but uses kernel_sysctl() instead. As a result,
ibcs2_sysconf() no longer needs Giant.
- The p_rlimit macro no longer exists.
Submitted by: mtm (mostly, I only did a few cleanups and catchups)
Tested on: i386
Compiled on: alpha, amd64
Presumably, at some point, you had to include jail.h if you included
proc.h, but that is no longer required.
Result of: self injury involving adding something to struct prison
setting the new process' p_pgrp again before inserting it in the p_pglist.
Without it we can get the new process to be inserted in a different p_pglist
than the one p2->p_pgrp points to, and this is not something we want to happen.
This is not a fix, merely a bandaid, but it will work until someone finds a
better way to do it.
Discussed with: jhb (a long time ago)
is useless for threaded programs, multiple threads can not share same
stack.
The alternative signal stack is private for thread, no lock is needed,
the orignal P_ALTSTACK is now moved into td_pflags and renamed to
TDP_ALTSTACK.
For single thread or Linux clone() based threaded program, there is no
semantic changed, because those programs only have one kernel thread
in every process.
Reviewed by: deischen, dfr
clobbers this variable. Long ago, when the idle loop wasn't in a
process, it set switchtime.tv_sec to zero to indicate that the time
needs to be read after the idle loop finishes. The special case for
this isn't needed now that there is an idle process (for each CPU).
The time is read in the normal way when the idle process is switched
away from. The seconds component of the time is only zero for the
first second after the uptime is set, and the mostly-dead code was only
executed during this time. (This was slightly broken by using uptimes
instead of times relative to the Epoch -- in the original version the
seconds component of the time was only 0 for the first second after
the Epoch.)
In mi_switch(), moved the setting of switchticks to just after the
first (and now only) setting of switchtime. This setting used to be
delayed since a late setting was needed for the idle case and an early
setting was not needed. Now the early setting is needed so that
fork_exit() doesn't need to set either switchtime or switchticks.
Removed now-completely-rotted comment attached to this. Most of the
code described by the comment had already moved to sched_switch().
begin with sched_lock held but not recursed, so this variable was
always 0.
Removed fixup of sched_lock.mtx_recurse after context switches in
sched_switch(). Context switches always end with this variable in the
same state that it began in, so there is no need to fix it up. Only
sched_lock.mtx_lock really needs a fixup.
Replaced fixup of sched_lock.mtx_recurse in fork_exit() by an assertion
that sched_lock is owned and not recursed after it is fixed up. This
assertion much match the one in mi_switch(), and if sched_lock were
recursed then a non-null fixup of sched_lock.mtx_recurse would probably
be needed again, unlike in sched_switch(), since fork_exit() doesn't
return to its caller in the normal way.
- Update some stale comments.
- Sort a couple of includes.
- Only set 'newcpu' in updatepri() if we use it.
- No functional changes.
Obtained from: bde (via an old diff I got a long time ago)
"maxproc limit exceeded by uid %i, please see tuning(7) and login.conf(5)."
Which will be triggered whenever a user hits his/her maxproc limit or
the systemwide maxproc limit is reached.
MFC after: 1 week
- Move struct sigacts out of the u-area and malloc() it using the
M_SUBPROC malloc bucket.
- Add a small sigacts_*() API for managing sigacts structures: sigacts_alloc(),
sigacts_free(), sigacts_copy(), sigacts_share(), and sigacts_shared().
- Remove the p_sigignore, p_sigacts, and p_sigcatch macros.
- Add a mutex to struct sigacts that protects all the members of the struct.
- Add sigacts locking.
- Remove Giant from nosys(), kill(), killpg(), and kern_sigaction() now
that sigacts is locked.
- Several in-kernel functions such as psignal(), tdsignal(), trapsignal(),
and thread_stopped() are now MP safe.
Reviewed by: arch@
Approved by: re (rwatson)
fini routines instead of in fork() and wait(). This has the nice side
benefit that the proc lock of any process on the allproc list is always
valid and sched_lock doesn't have to be used to test against PRS_NEW
anymore.
uptime. Where necessary, convert it back to Unix time by adding boottime
to it. This fixes a potential problem in the accounting code, which would
compute the elapsed time incorrectly if the Unix time was stepped during
the lifetime of the process.
rename them appropriately. Protect both flags with both the proc lock
and the sched_lock.
- Protect p_profthreads with the proc lock.
- Remove Giant from profil(2).
- Set p_acflag earlier while already hold the proc lock in fork1().
- Mark the realitexpire() callout MPSAFE for new processes. It was already
marked safe for proc0 a long while ago.
of ksegs since they primarily operation on processes.
- KSEs take ticks so pass the kse through sched_clock().
- Add a sched_class() routine that adjusts a ksegrp pri class.
- Define a sched_fork_{kse,thread,ksegrp} and sched_exit_{kse,thread,ksegrp}
that will be used to tell the scheduler about new instances of these
structures within the same process. These will be used by THR and KSE.
- Change sched_4bsd to reflect this API update.
if (p->p_numthreads > 1) and not a flag because action is only necessary
if there are other threads. The rest of the system has no need to
identify thr threaded processes.
- In kern_thread.c use thr_exit1() instead of thread_exit() if P_THREADED
is not set.
flexible process_fork, process_exec, and process_exit eventhandlers. This
reduces code duplication and also means that I don't have to go duplicate
the eventhandler locking three more times for each of at_fork, at_exec, and
at_exit.
Reviewed by: phk, jake, almost complete silence on arch@
struct proc as p_tracecred alongside the current cache of the vnode in
p_tracep. This credential is then used for all later ktrace operations on
this file rather than using the credential of the current thread at the
time of each ktrace event.
- Now that we have multiple ktrace-related items in struct proc that are
pointers, rename p_tracep to p_tracevp to make it less ambiguous.
Requested by: rwatson (1)
used to share resource limits between rfork threads, but never was.
Removing it makes resource limit locking much simpler -- only the current
process can change the contents of the structure that p_limit points to.
- Mark the process leader as having an advisory lock
- Check if process leader is marked as having advisory lock when
closing file
- Check that file is still open after lock has been obtained
- Don't allow file descriptor table sharing between processes
with different leaders
PR: 10265
Reviewed by: alfred
I'm not convinced there is anything major wrong with the patch but
them's the rules..
I am using my "David's mentor" hat to revert this as he's
offline for a while.
data structure called kse_upcall to manage UPCALL. All KSE binding
and loaning code are gone.
A thread owns an upcall can collect all completed syscall contexts in
its ksegrp, turn itself into UPCALL mode, and takes those contexts back
to userland. Any thread without upcall structure has to export their
contexts and exit at user boundary.
Any thread running in user mode owns an upcall structure, when it enters
kernel, if the kse mailbox's current thread pointer is not NULL, then
when the thread is blocked in kernel, a new UPCALL thread is created and
the upcall structure is transfered to the new UPCALL thread. if the kse
mailbox's current thread pointer is NULL, then when a thread is blocked
in kernel, no UPCALL thread will be created.
Each upcall always has an owner thread. Userland can remove an upcall by
calling kse_exit, when all upcalls in ksegrp are removed, the group is
atomatically shutdown. An upcall owner thread also exits when process is
in exiting state. when an owner thread exits, the upcall it owns is also
removed.
KSE is a pure scheduler entity. it represents a virtual cpu. when a thread
is running, it always has a KSE associated with it. scheduler is free to
assign a KSE to thread according thread priority, if thread priority is changed,
KSE can be moved from one thread to another.
When a ksegrp is created, there is always N KSEs created in the group. the
N is the number of physical cpu in the current system. This makes it is
possible that even an userland UTS is single CPU safe, threads in kernel still
can execute on different cpu in parallel. Userland calls kse_create to add more
upcall structures into ksegrp to increase concurrent in userland itself, kernel
is not restricted by number of upcalls userland provides.
The code hasn't been tested under SMP by author due to lack of hardware.
Reviewed by: julian
