Export limit exceeded: 26096 CVEs match your query. Please refine your search to export 10,000 CVEs or fewer.
Search
Search Results (26096 CVEs found)
| CVE | Vendors | Products | Updated | CVSS v3.1 |
|---|---|---|---|---|
| CVE-2026-99224 | 1 Linux | 2 Linux, Linux Kernel | 2026-11-02 | 7 High |
| In the Linux kernel, the following vulnerability has been resolved: keys: Protect the type name while describing a key keyctl_describe_key() must keep the key type's name alive until kasprintf() has consumed it. The retained key reference keeps the key and its description alive, but does not pin the module providing its type. The name is passed to kasprintf() without holding key->sem. kvasprintf() uses the saved name pointer in two formatting passes, with a GFP_KERNEL allocation between them. With AF_RXRPC=m, a task that passes View permission before type retirement can overlap a privileged module unload in this order: 1. KEYCTL_DESCRIBE looks up the key, saves key->type->name in the kasprintf() arguments and completes the first formatting pass. 2. While the formatter is delayed before its second pass, af_rxrpc_exit() calls unregister_key_type(), which removes the type from the registry and waits for key_gc_keytype(). 3. The collector takes key->sem for writing, retypes the retained key to key_type_dead and completes retirement. Unregistration and module exit return, allowing free_module() to release the type and name. 4. The formatter resumes its second pass and reads the saved name from released module storage, which can fault in string(). Hold key->sem for reading across kasprintf(). The collector then cannot retype the key or complete retirement until both formatting passes have finished. If retirement wins the lock first, formatting uses the core key_type_dead name instead. Release the semaphore immediately after kasprintf(), including on allocation failure, before copying to userspace. Oops report as below: BUG: unable to handle page fault for address: fffffbfff8054794 #PF: supervisor read access in kernel mode #PF: error_code(0x0000) - not-present page RIP: 0010:string+0x23f/0x470 Call Trace: vsnprintf+0x330/0x1110 | ||||
| CVE-2019-15291 | 1 Linux | 1 Linux Kernel | 2026-11-02 | N/A |
| An issue was discovered in the Linux kernel through 5.2.9. There is a NULL pointer dereference caused by a malicious USB device in the flexcop_usb_probe function in the drivers/media/usb/b2c2/flexcop-usb.c driver. | ||||
| CVE-2026-99220 | 1 Linux | 2 Linux, Linux Kernel | 2026-11-02 | N/A |
| In the Linux kernel, the following vulnerability has been resolved: asymmetric_keys: reject trust keys without IDs The PKCS#8 parser deliberately leaves the asym_key_ids payload slot empty. key_or_keyring_common() unconditionally dereferences that slot when a PKCS#8 key is supplied to a key_or_keyring restriction. Both the plain and :chain forms reach this branch, allowing an unprivileged caller to trigger a NULL pointer dereference through KEYCTL_RESTRICT_KEYRING followed by add_key(). Oops: general protection fault, probably for non-canonical address 0xdffffc0000000000 KASAN: null-ptr-deref in range [0x0000000000000000-0x0000000000000007] RIP: 0010:key_or_keyring_common (crypto/asymmetric_keys/restrict.c:205 crypto/asymmetric_keys/restrict.c:279) Call Trace: <TASK> __key_create_or_update (security/keys/key.c:884) key_create_or_update (security/keys/key.c:1021) __do_sys_add_key (security/keys/keyctl.c:134) do_syscall_64 (arch/x86/entry/common.c:52 arch/x86/entry/common.c:83) entry_SYSCALL_64_after_hwframe (arch/x86/entry/entry_64.S:130) </TASK> Kernel panic - not syncing: Fatal exception Reject an asymmetric restriction key without key IDs. A PKCS#8 private key cannot identify the signer of a candidate certificate and is not a valid trust anchor. | ||||
| CVE-2026-99222 | 1 Linux | 2 Linux, Linux Kernel | 2026-11-02 | N/A |
| In the Linux kernel, the following vulnerability has been resolved: keys: Protect the read method check against type removal keyctl_read_key() checks key->type->read without holding key->sem. A reference keeps the key alive, but does not pin the module containing its type. A surviving RxRPC key can therefore be read while AF_RXRPC is being unloaded: KEYCTL_READ task Module unload and key GC ---------------- ------------------------ lookup_user_key() permission / possession checks load key->type af_rxrpc_exit() unregister_key_type() GC: down_write(key->sem) GC: retype to key_type_dead GC: up_write(key->sem) unregister_key_type() returns free_module() read saved type's read member The final access can read freed module storage. The read invocation in __keyctl_read_key() is already protected by key->sem, but the earlier method check is outside that critical section. Move the method check into __keyctl_read_key(), after key_validate() and under its existing read hold on key->sem. The check and invocation then use the same protected type for length queries and every read retry. The GC cannot retype the key until the semaphore is released, and type unregistration waits for that GC pass before the module can be freed. Return -EOPNOTSUPP when validation succeeds and the type has no read method. A temporary buffer may now be allocated before this result; allocation and copying to userspace remain outside the semaphore. | ||||
| CVE-2026-99223 | 1 Linux | 2 Linux, Linux Kernel | 2026-11-02 | N/A |
| In the Linux kernel, the following vulnerability has been resolved: keys: make keyring indexes independent of kernel addresses KEYCTL_READ returns the serial numbers in a keyring in associative-array traversal order. The first word of the index is a public, deterministic hash of the type pointer, domain pointer and description. An unprivileged user can therefore insert keys with known descriptions, predict their order for each possible KASLR slide, and recover the kernel image address. On an x86-64 build with 16 MiB alignment, 45 user keys suffice to distinguish the candidate slides. Remove the addresses from the whole index, not just its first hash word. Assign each key type and domain a nonzero 64-bit ID and hash those IDs and the description with a once-initialized SipHash key. Use the full IDs in the later index chunks as well. Truncating or hashing the IDs there would lose the guarantee that distinct type/domain/description tuples have a differing index bit. Allocate IDs lazily since private key types need not be registered. Cache them in the index key so tree walks and copies of an existing index do not need to dereference a type descriptor that may have been unregistered. Keep the 32-bit chunk layout and diff_objects() bit offsets consistent, and retain the separate root slot for keyrings used by recursive search. The key comparison, permissions and userspace data format are unchanged; only the address-dependent in-memory ordering is replaced. | ||||
| CVE-2026-99230 | 1 Linux | 2 Linux, Linux Kernel | 2026-11-02 | N/A |
| In the Linux kernel, the following vulnerability has been resolved: io_uring/cmd: don't leave list poison in ->hash_node io_uring_cmd_del_cancelable() uses hlist_del() to take a command off ->cancelable_uring_cmd, which leaves LIST_POISON1/2 behind in ->hash_node. Requests are recycled through the request cache without being reinitialised, and not every user of ->hash_node sets it up before it is used: If a vectored futex wait is woken while futex_wait_multiple_setup() is still queueing the futexes, io_futexv_wait() doesn't add the request to ->futex_list, and leaves the completion to the task_work queued by the wake up. io_futexv_complete() ends in __io_futex_complete(), whose hlist_del_init() is fine for a request that is unhashed, which is how all other users of the node leave it. If the request was a cancelable uring_cmd in its previous life, e.g. a ublk FETCH_REQ, it writes through LIST_POISON2: Oops: general protection fault, probably for non-canonical address KASAN: maybe wild-memory-access in range [0xdead000000000120-...] RIP: 0010:io_futexv_complete+0x219/0x420 Call Trace: tctx_task_work_run+0x804/0x950 tctx_task_work+0x71/0xa0 task_work_run+0x118/0x200 Leave the node unhashed. | ||||
| CVE-2026-99231 | 1 Linux | 2 Linux, Linux Kernel | 2026-11-02 | 7.8 High |
| In the Linux kernel, the following vulnerability has been resolved: io_uring/waitid: own the request before clearing ->head io_waitid_remove_wq() only takes the wait queue lock if ->head is set. io_waitid_wait() clears ->head first thing, before it tries to become the owner of the request by bumping ->refs. If somebody else owns the request at that point, nothing keeps it from completing the request while the wake up handler is still going to touch it: CPU0 (io_waitid_wait) CPU1 (io_waitid, holds a ref) list_del_init(&wait->entry) smp_store_release(&iw->head, NULL) __do_wait() finds another child io_waitid_remove_wq() // ->head == NULL, no locking io_waitid_finish(), IOU_COMPLETE // request freed and reused atomic_fetch_inc(&iw->refs) The same goes for __io_waitid_cancel() if it claims the request in that window. ->refs is part of the per opcode data of the request, so this increments whatever the next user of the request keeps there. If that happens to be zero, the handler concludes that it owns the request, and queues io_waitid_cb() for it. With the window held open by a delay, and the request reused for a receive that is waiting for data: BUG: KASAN: slab-use-after-free in __wake_up_common+0x1ca/0x1f0 __wake_up_sync_key+0x36/0x50 sock_def_readable+0xf1/0x420 unix_stream_sendmsg+0x60b/0x1110 Allocated by task 87: io_arm_apoll+0x1ed/0x940 Freed by task 87: __io_submit_flush_completions+0x118d/0x1a10 tctx_task_work_run+0x327/0x950 Only clear ->head once the handler owns the request. If it doesn't, the owner finds ->head set and has to wait for the lock, which is held by the caller of the handler. Deleting the already deleted entry again is harmless. | ||||
| CVE-2026-99234 | 1 Linux | 2 Linux, Linux Kernel | 2026-11-02 | N/A |
| 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. | ||||
| CVE-2026-99228 | 1 Linux | 2 Linux, Linux Kernel | 2026-11-02 | N/A |
| In the Linux kernel, the following vulnerability has been resolved: io_uring/io-wq: don't lose the signal that ends work When a thread exits, only its tracked requests are cancelled. Then its io-wq is torn down. io_wq_exit_workers() signals every worker, to get it out of an interruptible sleep, and waits for all of them to exit. Before they do, they go through the work that is left. Unbounded work is cancelled. Bounded work is run, as it is expected to end by itself. Not all of it does. An open has no file yet, so it counts as bounded, and the open of a FIFO waits for the other end. That is what the signal is for. But the workers are signalled once, and a worker that hasn't started the work by then loses the signal: io_wq_exit_workers() io_worker_handle_work() io_wq_worker_wake() set TIF_NOTIFY_SIGNAL io_assign_current_work() io_run_task_work() clear TIF_NOTIFY_SIGNAL io_wq_submit_work() fifo_open() wait_for_partner() wait_for_completion_timeout() The thread never gets out of do_exit(): IORING_OP_OPENAT of a FIFO without a writer, IOSQE_ASYNC exit of the thread that has submitted it hangs in more than 9 of 10 attempts (4 CPUs, KVM), until the process is killed. Set TIF_NOTIFY_SIGNAL for bounded work that is started with IO_WQ_BIT_EXIT set. The bit is set before the workers are signalled, and it is tested after the flag has been cleared, with clear_notify_signal() providing the barrier. So either the flag stays, or the bit is seen. Work that is started while the io-wq exits is treated like work that is being run at that time: it isn't interrupted unless it waits interruptibly. | ||||
| CVE-2026-99217 | 1 Linux | 2 Linux, Linux Kernel | 2026-11-02 | N/A |
| In the Linux kernel, the following vulnerability has been resolved: fork: release the files of a failed fork after sched_cancel_fork() Reorder copy_process() cleanup so sched_fork() taking scx_fork_rwsem scx_pre_fork() is safe and isn't held around exiting files. Right now, a fork that fails while another thread closed the bpf link fd will deadlock against its own read side. It also blocks every fork on the system: copy_process() holds scx_fork_rwsem for read exit_files() close_files() filp_close_sync() bpf_scx_unreg() kthread_flush_work() waits for the disable work scx_root_disable() percpu_down_write(&scx_fork_rwsem) Simply release the child's files after sched_cancel_fork() dropped the lock. The task starts out as a copy of its parent so p->files points to the parent's table until copy_files() replaces it. exit_files() must not run for a fork that failed before that. Clear p->files up front so exit_files() is a no-op for those and make copy_files() set it explicitly for CLONE_FILES. | ||||
| CVE-2026-99218 | 1 Linux | 2 Linux, Linux Kernel | 2026-11-02 | 7 High |
| In the Linux kernel, the following vulnerability has been resolved: coredump: hold RCU while releasing parked threads coredump_finish() releases the parked threads by clearing ->task in their core_thread entry and calling wake_up_process() on each of them. It does that without holding a reference on the task and without being inside rcu. But calling wake_up_process(task) without rcu here isn't safe. A parked thread doesn't need that wakeup to leave. A spurious wakeup or a preemption after the store is enough for the task to go away. So if the coredump client is preempted between the store and wake_up_process() the thread can exit and be freed in the meantime and try_to_wake_up() takes pi_lock in freed memory. Hold rcu across the loop. | ||||
| CVE-2026-99187 | 1 Linux | 2 Linux, Linux Kernel | 2026-11-02 | 4.7 Medium |
| In the Linux kernel, the following vulnerability has been resolved: perf: Fix perf_pmu_migrate_context() leaving events behind __perf_pmu_remove() walks the group leaders of a CPU and PMU with perf_event_groups_for_cpu_pmu(), removing each one as it goes. perf_remove_from_context() takes the leader out of the groups rbtree, after which rb_next() on that node returns NULL, prematurely ending the walk. Other groups stay behind in the source context. On CPU hotplug, they are then torn down with the dying CPU. They keep running until teardown, while the PMU's interrupt and migrated events are on a new CPU, which drivers rely on never happening. Walk ctx->event_list instead and pick out the PMU's events by their pmu_ctx. | ||||
| CVE-2026-99189 | 1 Linux | 2 Linux, Linux Kernel | 2026-11-02 | 4.1 Medium |
| In the Linux kernel, the following vulnerability has been resolved: gpio: mxc: fix race between chained IRQ handler install and probe completion mxc_update_irq_chained_handler() is called before the IRQ domain, the generic IRQ chip, and the port list entry are set up. If an interrupt arrives in that window: - mx3_gpio_irq_handler() calls generic_handle_domain_irq() with port->domain still NULL. - mx2_gpio_irq_handler() walks mxc_gpio_ports, but the port has not been added to the list yet. Additionally, if any of the subsequent probe steps (gpio_generic_chip_init(), devm_gpiochip_add_data(), irq_domain_create_legacy(), or mxc_gpio_init_gc()) fail, the error paths never unregister the chained handler, leaving a dangling handler that points at freed memory. Move the handler installation after all its dependencies are ready and after list_add_tail(), so the handler is never live while the data structures it touches are incomplete, and is never installed if probe fails. | ||||
| CVE-2026-99181 | 1 Linux | 2 Linux, Linux Kernel | 2026-11-02 | 4.7 Medium |
| In the Linux kernel, the following vulnerability has been resolved: autofs: ignore notification errors from a replaced pipe autofs_notify_daemon() takes a reference to the current notification pipe under wq_mutex, then drops the mutex before writing the request. The write can block while the daemon enters catatonic mode and installs a replacement pipe. If the old pipe's reader is then closed, the delayed write returns -EPIPE and the generic error path enters catatonic mode again. That transition acts on the current superblock state, closes the replacement pipe, and releases the new daemon generation's wait queues. Only enter catatonic mode when the pipe which failed is still the current pipe. Factor the already-locked transition so the identity check and state change are atomic with respect to another replacement. This was reproduced on v7.3-rc5-337-gff47652a4b66c in a local QEMU guest. An unprivileged UID 65534 lookup filled the old packet pipe before a normal CATATONIC/SETPIPEFD restart. On the unmodified kernel, the new pipe received HUP and both a pending victim lookup and a later lookup failed with ENOENT in two out of two fresh-mount runs. With this change, the new pipe remained active, received an intact 304-byte request, and both victim lookups completed in two out of two runs. Normal lookup and restart-without-a-blocked-writer controls passed. No KASAN, oops, refcount, or lock diagnostic was observed. A source reproducer and complete local logs are available privately on request. The demonstrated impact is limited to denial of a shared autofs service during a legitimate daemon restart. No memory corruption, information leak, privilege escalation, or code execution primitive was observed. | ||||
| CVE-2026-99192 | 1 Linux | 2 Linux, Linux Kernel | 2026-11-02 | 4.7 Medium |
| In the Linux kernel, the following vulnerability has been resolved: net: bcmgenet: unshare the skb before writing the control block bcmgenet stores first_cb, last_cb and bytes_sent in the skb control block and reads them back on completion to free the skb and credit BQL. ether_setup() leaves IFF_TX_SKB_SHARING set, so the stack can resubmit a shared skb that is still in the ring. The next transmit overwrites last_cb, so the earlier slots complete without crediting their bytes and BQL stalls the queue until the watchdog resets it. Unshare the skb before writing the control block. Reproduce with pktgen in clone_skb mode, which hands the driver the same skb while earlier copies are still in the ring: modprobe pktgen pg=/proc/net/pktgen echo "add_device eth0" > $pg/kpktgend_0 echo "dst_mac AA:BB:CC:DD:EE:FF" > $pg/eth0 echo "dst 192.0.2.1" > $pg/eth0 echo "clone_skb 8" > $pg/eth0 echo "count 0" > $pg/eth0 echo start > $pg/pgctrl & sleep 5 echo stop > $pg/pgctrl wait cat /sys/class/net/eth0/queues/tx-0/byte_queue_limits/inflight inflight stays nonzero on the now idle queue. clone_skb 0 leaves it at zero. | ||||
| CVE-2026-99202 | 1 Linux | 2 Linux, Linux Kernel | 2026-11-02 | 4.1 Medium |
| In the Linux kernel, the following vulnerability has been resolved: scsi: libsas: Fix SMP IO deadlock during HA resume smp_execute_task_sg() calls pm_runtime_get_sync() on the host before issuing an SMP command. When that command is itself issued from the HA resume path, the get_sync() deadlocks: it waits for the ongoing resume (the device is RPM_RESUMING), while the resume is blocked in sas_drain_work() waiting for that same SMP IO to complete. The deadlock needs an expander-attached SATA disk. During sas_resume_ha() -> sas_drain_work(), DISCE_RESUME -> sas_resume_sata() -> ata_sas_port_resume() requests ATA_EH_RESET, and the hard reset for such a disk is done via SMP PHY CONTROL (sas_ata_hard_reset() -> sas_phy_reset() -> sas_smp_phy_control() -> smp_execute_task_sg()). Direct-attached SATA resets through lldd_control_phy() and SSP devices use TMFs, so neither hits this. Replace the get_sync()/put_sync() pair with pm_runtime_get_noresume()/pm_runtime_put(). smp_execute_task_sg() only needs to hold off autosuspend while the SMP is in flight, and it must not try to resume the host: a sync resume issued from the HA resume path itself is what deadlocks, and by the time sas_resume_ha() runs, hw_init has already reinitialized the hardware, so the device is accessible without one. The usage reference is still required. Discovery work normally runs inside an event worker's PM reference, taken at sas_notify_port_event() notify time and held until the handler has flushed the disco queue. sas_rediscover_ex_phy() however requeues DISCE_REVALIDATE_DOMAIN from within the revalidation worker itself, and flush_workqueue() does not wait for work items queued during execution, so that chained revalidation runs with no outer PM reference - without the get_noresume(), its SMP could race autosuspend. For the BSG path, sas_smp_handler() is the only caller which may find the host autosuspended: expander SMP requests do not go through any SCSI device request queue, so nothing else in that path holds the host awake. Resume it there with pm_runtime_resume_and_get() and check the result. | ||||
| CVE-2026-99159 | 1 Linux | 2 Linux, Linux Kernel | 2026-11-02 | 4.4 Medium |
| In the Linux kernel, the following vulnerability has been resolved: ring-buffer: Fix the sub-buffer array base in rb_meta_subbuf_idx() On a persistent ring buffer, rb_meta_subbuf_idx() can return an index one sub-buffer too high. rb_setup_ids_meta_page() then fills subbuf_ids[] shifted by one and leaves subbuf_ids[0] NULL, which __rb_map_vma() dereferences when userspace maps trace_pipe_raw. rb_update_meta_reader() writes two such indices into meta->buffers[], which rb_cpu_meta_valid() would reject on the next boot, discarding the previous boot's trace. A CPU's sub-buffer array starts after its meta header and the nr_subbufs integers that follow, aligned up to a sub-buffer. rb_meta_subbuf_idx() inverts that calculation but skips only the integers, not the header, so its base is one sub-buffer low whenever the omitted header size crosses the alignment boundary. Invert the calculation with rb_subbufs_from_meta() rather than repeating it. rb_range_buffer() maps an index back to an address through the same helper, so the two directions can no longer disagree. | ||||
| CVE-2026-99154 | 1 Linux | 2 Linux, Linux Kernel | 2026-11-02 | 9.8 Critical |
| In the Linux kernel, the following vulnerability has been resolved: scsi: libiscsi: Reject a negative task index from parse_pdu_itt A transport that implements parse_pdu_itt hands libiscsi an index taken from the PDU, and both lookups bound it from above only: if (i >= session->cmds_max) i and cmds_max are both int, so a negative index passes and session->cmds[i] is read from before the array. be2iscsi produces one. beiscsi_parse_pdu() assigns the raw tag, *index = (int)itt; and beiscsi_complete_pdu() forwards an unsolicited NOP-In from the hardware async ring without replacing its ITT, so the value is the target's. It also reports the session's own age rather than the one in the tag, which leaves the age comparison in iscsi_verify_itt() with nothing to reject. Bound the index from below in both lookups. | ||||
| CVE-2026-99156 | 1 Linux | 2 Linux, Linux Kernel | 2026-11-02 | 7.8 High |
| In the Linux kernel, the following vulnerability has been resolved: scsi: st: Fix use-after-free race between st_remove() and st_open() st_remove() drops the probe reference on the scsi_tape with kref_put() and only afterwards removes its index from st_index_idr. If that was the last reference, scsi_tape_release() frees the structure while it is still published in the idr. scsi_tape_get(), called from st_open(), looks up the tape in st_index_idr and takes a reference under st_ref_mutex and st_index_lock. remove_cdevs() does not wait for opens that are already past chrdev_open(), so an open that is blocked on st_ref_mutex while st_remove() holds it can take the mutex as soon as st_remove() drops it, find the freed tpnt in the idr and call kref_get() on it: BUG: KASAN: slab-use-after-free in st_open Write of size 4 ... st_open (drivers/scsi/st.c:254 drivers/scsi/st.c:1291) chrdev_open ... Freed by task ...: kfree st_remove (drivers/scsi/st.c:4512) Remove the index from the idr before dropping the reference, all under st_ref_mutex. A concurrent scsi_tape_get() now either runs first and holds a valid reference that keeps the tape alive, or runs afterwards and finds nothing, so st_open() fails with -ENXIO. This restores the ordering st_remove() had before the conversion to st_index_idr, where the tape was unpublished from scsi_tapes[] before the final kref_put(). | ||||
| CVE-2026-99158 | 1 Linux | 2 Linux, Linux Kernel | 2026-11-02 | 4.1 Medium |
| In the Linux kernel, the following vulnerability has been resolved: ring-buffer: Fix false warning in ring_buffer_map_get_reader() The mmap reader can warn when it catches up with a writer that is still committing events. rb_get_reader_page() returns NULL in this case, but ring_buffer_map_get_reader() treats that result as an error. The initial rb_per_cpu_empty() check filters out an empty buffer, but it does not guarantee that the next reader page has committed data. Writers do not take reader_lock and can advance commit_page before publishing the committed length on the new page. rb_get_reader_page() can swap to that page, catch up with the writer and return NULL. Handle NULL through the existing no-data exit without warning. Remove the caller's reader_page == commit_page check as well, since the reader-page helper already handles that case. The existing metadata update and return path remain in place. | ||||