| CVE |
Vendors |
Products |
Updated |
CVSS v3.1 |
| In the Linux kernel, the following vulnerability has been resolved:
io_uring/napi: make stale entries time out after 60s
With dynamic tracking a NAPI id that hasn't been seen for a minute is
meant to be dropped from the list the ring busy polls. The timeout is
added to jiffies, but it is defined as 60 * SEC_CONVERSION, and
SEC_CONVERSION isn't the number of jiffies in a second. It is the
internal factor of __timespec64_to_jiffies(), HZ scaled up by
SEC_JIFFIE_SC bits. For HZ=1000 that is 1000 << 22, which puts the
timeout at about eight years. On 32-bit the multiplication overflows.
So entries are never removed, and every id that a socket of the ring
ever had keeps being polled on each wait. |
| In the Linux kernel, the following vulnerability has been resolved:
io_uring/zcrx: fix buffer offsets in areas above 4GB
io_zcrx_queue_cqe() computes the offset of the buffer in its area from
net_iov_idx(), which is an unsigned int, so the shift is done in 32 bits
and only the result is widened to the u64 that goes into the CQE. An area
can be up to 1TB in size. For every niov beyond the first 4GB the CQE
points the application at a buffer 4GB, or a multiple of that, below the
one which has the data. It reads the wrong data, and then returns a
buffer it doesn't own via the refill queue while the real one is never
given back. |
| In the Linux kernel, the following vulnerability has been resolved:
io_uring/openclose: keep the lock in io_pipe_fixed()
io_pipe_fixed() installs both ends of the pipe in the fixed file table,
drops ->uring_lock, and copies the slot numbers to the application. If
that faults, it takes the lock again and removes whatever is in these
slots.
If the request is run by io-wq the lock really is dropped in between, and
the slots are visible to everybody else. By the time the error path gets
to them, they may have been closed and reused, by IORING_OP_CLOSE and a
direct open or accept, or by IORING_REGISTER_FILES_UPDATE. Then a file is
removed from the table that has nothing to do with the failed request.
Do the copy with the lock held. Faulting in user memory under
->uring_lock is fine, the registration paths do it all the time. |
| In the Linux kernel, the following vulnerability has been resolved:
io_uring/poll: avoid retry wakeup loss when dropping ownership
If a wakeup sets IO_POLL_RETRY_FLAG between the owning task consuming
the existing instance of IO_POLL_RETRY_FLAG and the loop check
of io_poll_check_events(), then we stall on processing the request
until an unrelated wakeup comes along.
Using atomic_cmpxchg instead of atomic_sub_return to guarantee both the
count v and poll_refs are 0 and no wake-ups are missed. |
| In the Linux kernel, the following vulnerability has been resolved:
io_uring/timeout: don't lose a cancel as io-wq arms a timeout
An IORING_OP_TIMEOUT run by io-wq, because of IOSQE_ASYNC, has the same
problem as poll, futex and waitid. io_wq_submit_work() looks at
IO_WQ_WORK_CANCEL once, before it issues the request. A cancel in between
flags the work, gets -EALREADY, and searches ->timeout_list in vain,
while the worker then adds the request and starts the timer:
io-wq worker io_try_cancel()
io_wq_submit_work()
IO_WQ_WORK_CANCEL isn't set
io_async_cancel_one()
sets IO_WQ_WORK_CANCEL
= -EALREADY
io_timeout_cancel() = -ENOENT
io_timeout()
adds the request to ->timeout_list
hrtimer_start()
= IOU_ISSUE_SKIP_COMPLETE
Nobody looks at the flag again, and the timeout runs its course. At ring
exit, io_uring_try_cancel_requests() cancels io-wq work and then kills
the timeouts the same way, so an exiting task can get stuck for the full
duration of such a timeout.
Look at the flag under ->timeout_lock, which the search takes as well,
and fail the request if it is set. If the search came first, the flag is
visible. If it comes later, it finds the request on the list. |
| In the Linux kernel, the following vulnerability has been resolved:
io_uring/io-wq: don't hide a link from cancellations
A task that exits cancels its tracked requests and waits until they are
gone. io_uring_cancel_generic() looks everywhere a request can be, and
goes to sleep if it found nothing to cancel. It is woken up each time
one of the requests of the task is completed by task work or freed.
When io-wq is done with a request that has a link, there is a time in which
the next request of the link can't be found:
io_worker_handle_work()
io_assign_current_work(worker, NULL)
io_wq_free_work(work)
io_req_find_next() // takes the link off the request
io_free_req() // task work for the task
io_assign_current_work(worker, linked)
io_run_task_work()
cond_resched()
worker->cur_work = linked
It is not pending, it is not the current work of a worker, and it is not
yet on the list that it is about to be parked on. If the pass that the
exiting task makes because of io_free_req() falls into that time, it finds
nothing and the task goes to sleep. The worker then runs the link, which
IO_WQ_BIT_EXIT doesn't prevent for bounded work. If that parks the request
nothing is completed or freed, and nobody wakes up the task. If it is the
process that exits it can't be killed any more either.
NOP, IOSQE_ASYNC | IOSQE_IO_LINK
FUTEX_WAIT, FUTEX2_PRIVATE
exit of the thread, or of the process, up to 63us later
With four of these per round 5 to 8 of 10,000 rounds hang (4 CPUs, KVM):
io_uring_cancel_generic
do_exit
Commit f8b632e89a10 ("io_uring: tighten task exit cancellations") is about
the same gap, and makes the task look again for whatever happens before
io_free_req().
Free the request once the link is the current work of the worker, or
pending. Whoever is woken up by that can find the link, until it is parked
where cancellations look as well. To do so split io_wq_free_work() in two. |
| In the Linux kernel, the following vulnerability has been resolved:
io_uring: don't flood a cancelling task with task_work
io_ring_exit_work() asks each task which has used the ring to remove its
tctx node, by task_work, and repeats that until the list is empty.
io_tctx_exit_cb() doesn't touch the node while the task is in
io_uring_cancel_generic(), which removes all the nodes at its end. So the
work item queues the callback again as soon as it has completed, for
nothing.
That is more than a waste of time. task_work_run() only returns when it
finds the list empty. If the rings are released from inside the
cancellation, because requests of other rings held the last references to
their files, up to 64 work items keep sending callbacks to the one task
which is running them. It never finds the list empty, never gets back to
the cancellation, the nodes stay, and so it goes on: a task with 60 pairs
of rings which poll each other ran 280000 callbacks in 5 seconds, 20
seconds after exit_group(). It ends when all the senders happen to be late
at the same time. Until then the task keeps a CPU and the work items busy.
An exiting task can't be killed meanwhile. execve() cancels the requests in
the same way, so a live task can get stuck there as well, for minutes. If
it takes more than five, the work items run into the
WARN_ON_ONCE(time_after(jiffies, timeout)) of io_ring_exit_work(). No
privilege is needed for any of this.
If the callback has met a task in cancellation, wait a little before it is
sent again. Nobody has to wake the work item up for that: either the node
is gone by then, or it is the turn of the next round. |
| In the Linux kernel, the following vulnerability has been resolved:
io_uring/cancel: rearm local work wake up before sleep
io_uring_try_cancel_requests() sets ->cq_wait_nr to 1 so that the next
io_req_local_work_add() wakes the cancelling task. The add that takes the
count to zero is the only one that wakes, later ones see a count <= 0 and
return. io_uring_cancel_generic() however only becomes
TASK_INTERRUPTIBLE much later, so the one wake up can be spent on a
running task. That is fine as long as the entry of that add is still
queued, as the io_local_work_pending() check then skips the sleep. It is
not if the add was delayed after queueing its entry:
CPU0 (submitter in exec or exit) CPU1 (io_req_local_work_add)
mpscq_push(req A)
io_uring_try_cancel_requests()
io_run_local_work() // runs req A
io_uring_try_cancel_requests()
atomic_set(&ctx->cq_wait_nr, 1)
atomic_read(&ctx->cq_wait_nr) == 1
count taken, wake_up_state()
// task is running, no-op
prepare_to_wait(TASK_INTERRUPTIBLE)
io_local_work_pending() == false
schedule()
mpscq_push(req B)
cq_wait_nr <= 0, return
Only the sleeping task can run req B, so it is stuck in execve() until a
signal happens to wake it. An exiting task doesn't even get that.
Arm the count again once the task state is set and before the work list
is checked, for the rings whose local work this task is the one to run. |
| In the Linux kernel, the following vulnerability has been resolved:
io_uring/mpscq: disable preemption while pushing
A push makes the node the new tail, and then links it to its predecessor.
In between the queue isn't empty, but mpscq_pop() can't get at the node.
The consumers deal with that by trying again, with a cond_resched() in
the loop to give a preempted producer the chance to finish.
cond_resched() doesn't do that if the consumer has a higher priority than
the producer and both have to use the same CPU:
producer (SCHED_OTHER) consumer (SCHED_FIFO), same CPU
io_req_local_work_add()
mpscq_push()
prev = xchg(&q->tail, node)
<preempted> io_cqring_wait()
__io_run_local_work()
again:
cond_resched() // nothing to do
mpscq_pop() // NULL
// work pending, no events yet
goto again
The producer can be any task that is able to queue task_work for the
ring: an io-wq worker, or somebody who sends a message with
IORING_OP_MSG_RING. The consumer spins, holding ->uring_lock, until the
scheduler takes pity on the fair class, which takes about a second with
the default settings of the fair server, and forever without it. The
loops in tctx_task_work_run() and io_cancel_local_task_work() have the
same problem.
Don't allow preemption between the two steps. That bounds the time the
window is open on the CPU of the producer by what interrupts take, and a
consumer that waits for that is on another CPU. The cost is a
preempt_disable()/preempt_enable() pair per push. |
| In the Linux kernel, the following vulnerability has been resolved:
KEYS: Fix key_reject_and_link() race with keyring restriction
The restrict_link check in key_reject_and_link() is done before taking
the keyring semaphore, racing with a concurrent keyring_restrict()
that installs the restriction while holding the semaphore in write
mode.
Move the check to after __key_link_lock(), so it is done under the
semaphore, matching the pattern used in __key_create_or_update(). |
| In the Linux kernel, the following vulnerability has been resolved:
pid_namespace: prevent TIF_NOTIFY_SIGNAL from interrupting the reaper
Since commit 7fea700e04bd ("zap_pid_ns_processes: clear
TIF_NOTIFY_SIGNAL along with TIF_SIGPENDING") zap_pid_ns_processes()
clears TIF_NOTIFY_SIGNAL before every kernel_wait4() call. Otherwise
kernel_wait4() doesn't sleep and returns -ERESTARTSYS if
signal_pending() is true and the reaper busy-waits.
But clearing TIF_NOTIFY_SIGNAL doesn't fix it. It can get raised again.
For example, klp_send_signals() sets it raw against any task that hasn't
transitioned yet. And it keeps sending it until the livepatch finishes.
The sleep loop waiting for pid_allocated to drop has the same problem
and it doesn't even clear anything. So a schedule() in
TASK_INTERRUPTIBLE returns immediately.
Use a no_notify_signal guard. signal_pending() will ignore
TIF_NOTIFY_SIGNAL and the pid namespace reaper sleeps until a child
exits or free_pid() wakes it. |
| In the Linux kernel, the following vulnerability has been resolved:
io-wq: order the exit bit against worker creation task work
io_wq_exit_workers() cancels task work for queued worker creation with
task_work_cancel_match(). That has a plain load of task->task_works.
io_queue_worker_create() adds the new entry and tests IO_WQ_BIT_EXIT to
cancel it if the workqueue is already on its way out.
That must be ordered. task_work_add() has a full barrier via cmpxchg().
But io_wq_exit_start() sets the bit with set_bit() which doesn't have
any memory ordering.
So afaict, on weakly ordered architectures the exiting task may load
task_works before the store of the bit is visible. The other side tests
the bit before the store is visible as well. So the entry remains queued
with worker_refs and the exiting task waits on worker_done indefinitely.
If the task still has rings then io_uring_del_tctx_node() provides the
barrier via test_and_set_bit() in io_wq_set_exit_on_idle(). When it
has closed all rings though that barrier is gone.
Add the barrier after the set_bit(). |
| In the Linux kernel, the following vulnerability has been resolved:
Bluetooth: L2CAP: refuse __l2cap_chan_add if chan already has conn
l2cap_chan may be linked to l2cap_conn at most once. This is assumed in
several places, eg l2cap_chan_del cleanup.
There is a TOCTOU race where the invariant is violated:
[Task 1] [Task 2]
l2cap_chan_connect l2cap_sock_bind
l2cap_chan_lock lock_sock
l2cap_state_change if (sk->sk_state != BT_OPEN)
chan->state = BT_CONNECT
l2cap_sock_state_change_cb chan->state = BT_BOUND
sk->sk_state = BT_BOUND
lock_sock <------------------ release_sock
sk->sk_state = BT_CONNECT
l2cap_sock_connect() does not check sk->sk_state, so since chan->state
is now BT_BOUND, subsequent connect() ends up with second
__l2cap_chan_add.
Explicitly document and check the invariant in __l2cap_chan_add with
WARN_ON_ONCE. The only callsite where it could be hit is
l2cap_chan_connect, so add pre-check there to avoid relying on
chan->state. chan->state read/write is not properly guarded currently so
there can be other TOCTOUC problems.
Add l2cap_lock_chan in l2cap_sock_bind() to guard chan->state write. |
| In the Linux kernel, the following vulnerability has been resolved:
tracing: Hold trace_types_lock when updating modules in persistent instances
trace_module_record() walks ftrace_trace_arrays from the module notifier
without trace_types_lock. A concurrent rmdir of an instance removes it
from the list and frees it under that lock, so a module load or unload
can walk into a freed trace_array. |
| In the Linux kernel, the following vulnerability has been resolved:
tracing: Hold trace_types_lock when resetting last boot data from readers
update_last_data_if_empty() is called from tracing_open() with O_TRUNC,
tracing_read_pipe() and tracing_buffers_read() without trace_types_lock.
Two readers can both pass the LAST_BOOT test and free module_delta
twice, and a module notifier can update module_delta while it is being
freed.
Take trace_types_lock in those paths and keep the callers that already
hold it on an unlocked variant. |
| In the Linux kernel, the following vulnerability has been resolved:
tracing: Reject incomplete event filter expressions
A non-empty expression normally finishes in predicate_parse()'s inner
loop, which resolves the pending logical operators before jumping to
finalization. If the input ends after an operator, possibly followed by
unary '!' tokens, the outer loop instead falls through with unresolved
operator state.
For example, "common_pid == 1 || common_pid == 2 &&" leaves
prog[0].target equal to -1. The second pass then reads prog[-1] before
the forward-jump check can reject it. Depending on adjacent slab
contents, it can also copy an out-of-bounds word into a branch target
that passes the forward-only check. A self-targeted perf tracepoint
counting event can reach this path through PERF_EVENT_IOC_SET_FILTER.
Reject the outer-loop fallthrough when a predicate has already been
parsed. This ensures that non-empty expressions only reach finalization
after the inner loop has resolved their operators. Keep the existing
no-filter error for input that has no predicates. Add a regression test
for trailing logical operators and unary negations. |
| In the Linux kernel, the following vulnerability has been resolved:
tracing: Save the event file instead of the compatible histogram
create_field_var_hist() creates a histogram on the matched event to
supply a field variable, and records it on the target so it can be torn
down later:
hist_data = find_compatible_hist(target_hist_data, file);
...
/* Save the compatible histogram information */
var_hist->hist_data = hist_data;
hist_data is not the histogram it creates. It is one that was already
there, whose keys the new one copies. The saved pointer has a single
user, unregister_field_var_hists(), which runs when the target is
removed:
file = hist_data->field_var_hists[i]->hist_data->event_file;
find_compatible_hist() matches on keys alone, so it can pick one with no
variables of its own. check_var_refs() then returns false and the user
can remove it while the target still points at it:
BUG: KASAN: slab-use-after-free in event_hist_trigger_free+0x2b2/0x320
Read of size 8 at addr ffff8880093b90e0 by task init/1
Freed by task 1:
kfree+0x154/0x420
event_hist_trigger_free+0x1d5/0x320
event_hist_trigger_parse+0x35f3/0x69e0
trigger_process_regex+0x1a6/0x250
Save the event file instead. It is all the dereference ever wanted, and
it does not go away when the user removes a trigger. |
| In the Linux kernel, the following vulnerability has been resolved:
tracing: Unregister field variable histograms when the command fails
An action that names a field on another event needs that field as a
variable, so create_field_var_hist() registers a second hist trigger on
that event to provide it, and records it in
target_hist_data->field_var_hists[].
If a later part of the same command fails, create_actions() returns an
error and event_hist_trigger_parse() jumps to out_free. The write fails
with -EINVAL and sched_switch gets no trigger, but:
# echo 'hist:keys=pid:ts0=common_timestamp.usecs' >> \
events/sched/sched_waking/trigger
# echo 'my_synth u64 lat; int a; int b' >> synthetic_events
# echo 'hist:keys=next_pid:wakeup_lat=common_timestamp.usecs-$ts0:\
onmatch(sched.sched_waking).my_synth($wakeup_lat,prio,nosuchfld)'\
>> events/sched/sched_switch/trigger
# cat events/sched/sched_waking/trigger
hist:keys=pid:vals=hitcount:ts0=common_timestamp.usecs:sort=hitcount:size=2048:clock=global [active]
hist:keys=pid:vals=hitcount:synthetic_prio=prio:sort=hitcount:size=2048 [active]
The second histogram is the one created for 'prio'. It is still active and
still counting.
out_free reaches destroy_field_var_hists(), which frees the
field_var_hists[] entries, but never unregister_field_var_hists(), which is
what removes the triggers and runs only from event_hist_trigger_free(). The
entry is freed and the trigger is left with nothing tracking it.
Unregister the field variable histograms in out_free as well, and record
each one as soon as its trigger is registered rather than after the lookup
that follows it, so no exit can leave a registered trigger untracked. The
entry then belongs to field_var_hists[] and is released by
destroy_field_var_hists(), so freeing it on that exit would be a double
free. |
| In the Linux kernel, the following vulnerability has been resolved:
tracing/histograms: Verify event_call identity when updating synth_event refcount
In event_hist_trigger_parse() and hist_unreg_all(), the histogram
registration and unregistration paths look up find_synth_event() using
trace_event_name(file->event_call) and adjust se->ref if a match is
found.
Because find_synth_event() only compares the event basename, registering
or removing a histogram on a non-synthetic trace event whose basename
matches a synthetic event erroneously modifies the synthetic event's
reference count. If a histogram is registered on the non-synthetic
event before the synthetic event is created and removed afterwards,
se->ref is decremented without a matching increment, allowing the
synthetic event to be freed while an active histogram action still holds
a pointer to it.
Introduce find_synth_event_call() to verify that &se->call matches
file->event_call before adjusting se->ref. |
| In the Linux kernel, the following vulnerability has been resolved:
tracing/syscalls: Use copy_from_user for the faultable syscall buffer
syscall_copy_user_array() copies the buffer argument of faultable
syscalls (write, pwrite64, ...) with __copy_from_user(), which skips
access_ok(). Use copy_from_user() instead so a non-user address is
rejected and the entry is recorded as faulted.
It could be a security concern for machines that trust perf a bit too
much. |
