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Search Results (408480 CVEs found)
| CVE | Vendors | Products | Updated | CVSS v3.1 |
|---|---|---|---|---|
| CVE-2026-93251 | 1 Linux | 1 Linux Kernel | 2026-10-15 | N/A |
| In the Linux kernel, the following vulnerability has been resolved: ACPI: bus: Introduce acpi_bus_get_primary_device() The function used for obtaining the first "physical" device for which the given ACPI one is the ACPI companion, acpi_get_first_physical_node(), may return a stale device pointer (mostly in theory) because acpi_unbind_one() may run as a whole after dropping the ACPI device's physical_node_lock in acpi_get_first_physical_node() and before it returns. The last reference to the "physical" device may be dropped then before the pointer to it is returned to the caller. If that happens and the acpi_get_first_physical_node() caller invokes get_device() on the pointer obtained from it, which is done by the majority of its callers, a use-after-free will occur. To prepare for addressing this problem, introduce a new function for getting the first "physical" device associated with the given ACPI one (the "primary physical device") that will also reference count the device in question before returning a pointer to it. Make that new function and acpi_get_first_physical_node() share the physical node list lookup code. No intentional functional impact. | ||||
| CVE-2026-93248 | 1 Linux | 1 Linux Kernel | 2026-10-15 | N/A |
| In the Linux kernel, the following vulnerability has been resolved: drm/xe: don't WARN on kernel job timeout when device already wedged igt@xe_wedged@wedged-at-any-timeout wedges the device in mode 2 (UPON_ANY_HANG_NO_RESET) and then rebinds the driver. During unbind, a GSC proxy kernel submission can still time out; with the device wedged and the GuC CT stopped it can never complete, so its kernel job times out. Tile0: GT1: Kernel-submitted job timed out WARNING: drivers/gpu/drm/xe/xe_guc_submit.c:... at guc_exec_queue_timedout_job() Workqueue: gt-ordered-wq drm_sched_job_timedout Killed queues skip guc_submit_hint_wedged(), leaving 'wedged' false even though the device is already wedged. The timeout handler then treats the kernel queue timeout as unexpected and taints the kernel. Honour an already-wedged device even for killed queues so the expected teardown timeout no longer trips the WARN. (cherry picked from commit a1c1dbd0f047bb05de6aaf6abe9103031179bf19) | ||||
| CVE-2026-93243 | 1 Linux | 1 Linux Kernel | 2026-10-15 | N/A |
| In the Linux kernel, the following vulnerability has been resolved: mm/secretmem: properly account locked pages secretmem accounts folios by treating memory as if it were mlock()'d and thus limited by the RLIMIT_MEMLOCK limit. However the folios are unevictable and remain so until the inode is evicted, eliminating usual mlock() semantics - mapping folios then unmapping them does not clear their unevictable state, since it depends on AS_UNEVICTABLE, not PG_mlocked. A user can therefore easily work around the RLIMIT_MEMLOCK limit - simply map then unmap and VmLck no longer counts the secretmem range. Worse, folios are not accounted in the process's RSS, meaning the OOM killer won't know to kill the process. Repeatedly mapping/unmapping (or forking) can then result in the consumption of all available system memory with unevictable folios and cause system instability. A secretmem fd can be passed between processes and over fork so a per-process limit simply does not make sense, so follow the precedent set by io_uring, perf, skbuff, iommufd and xdp by tracking the number of locked pages in user_struct->locked_vm. Since the scope tracked is actually inode lifetime, the RLIMIT_MEMLOCK applies per-user not per-process, so it doesn't make sense to bypass for users with CAP_IPC_LOCK, therefore remove this bypass. There is simply no reason to carry on marking the mapping as mlock()'d since it's misleading and the lifecycle is now correctly handled, so remove this too. Note that secretmem does not support any form of truncation (including hole punching) and the folios are unreclaimable, so the folios need only be accounted on fault and unaccounted on inode destruction. __secretmem_account_pages() is more or less a duplicate of the code that io_uring etc. use, but since this is a bug fix that needs backporting, defer any de-duplication efforts to a follow-up. test_mlock_limit() asserts mlock_future_ok() on mmap(), however this has been removed, so remove the test altogether for the fix. A new test will be sent separately for upstream. | ||||
| CVE-2026-93221 | 1 Linux | 1 Linux Kernel | 2026-10-15 | 8.1 High |
| In the Linux kernel, the following vulnerability has been resolved: nfsd: convert nfsd_net boolean flags to unsigned long flags word nfsd_net contains several boolean fields that are accessed from concurrent contexts without serialization. In particular, nfsd4_end_grace() guards its drain path with a plain bool: if (nn->grace_ended) return; nn->grace_ended = true; The read and the write are independent, and nothing in struct nfsd_net serializes them. At least two contexts can reach this code with no lock held: laundromat path laundry_wq kworker nfs4_laundromat() nfsd4_end_grace() RECLAIM_COMPLETE path nfsd compound kthread nfsd4_reclaim_complete() inc_reclaim_complete() nfsd4_end_grace() Both callers can observe grace_ended == false on different CPUs, both store true, and both proceed into nfsd4_record_grace_done(), which invokes the active client_tracking_ops->grace_done callback. For tracking ops that drain reclaim_str_hashtbl (legacy_tracking_ops via nfsd4_recdir_purge_old, and the cld v1+ ops via nfsd4_cld_grace_done), grace_done calls nfs4_release_reclaim(), which walks every bucket of reclaim_str_hashtbl with no lock and calls nfs4_remove_reclaim_record() (list_del + kfree) on each entry. Two concurrent walkers corrupt the list and double-free every nfs4_client_reclaim. A concurrent nfsd4_find_reclaim_client() iterating the same bucket reads through freed memory. A third call site exists in nfs4_state_start_net() on the skip_grace startup path, but it runs under nfsd_mutex before any client has connected and before the laundromat's first delayed work fires, so it cannot race with the two callers above. Replace the scattered boolean fields in nfsd_net with a single unsigned long flags word and an enum nfsd_net_flag for the bit positions. The grace_ended race is fixed by using test_and_set_bit(), which is atomic on all architectures. The remaining flags (grace_end_forced, in_grace, somebody_reclaimed, track_reclaim_completes, nfsd_net_up, lockd_up) are converted to use test_bit/set_bit/clear_bit for consistency. This avoids sub-word cmpxchg issues on architectures like Hexagon that only support word-sized atomic operations. | ||||
| CVE-2026-93201 | 1 Linux | 1 Linux Kernel | 2026-10-15 | 7.8 High |
| In the Linux kernel, the following vulnerability has been resolved: dm-pcache: validate seg_id fields from persistent memory cache_pos_decode(), cache_key_decode() and the last-kset branches of cache_replay(), the writeback worker and the GC worker take a cache segment id from the cache device metadata and index cache->segments[] with it without checking it against cache->n_segs. That metadata is only CRC-protected with a fixed public seed, so whoever supplies the cache device on a table load (CAP_SYS_ADMIN) controls the id; an out-of-range value forms a wild pcache_cache_segment pointer that is dereferenced and written through -- an out-of-bounds read and write driven by on-disk data. Add cache_seg_id_valid() and reject an out-of-range id at each decode site, failing the operation with -EIO instead of indexing past the array. Bound the id against the initialized-segment count (cache_info.n_segs) rather than the physical device total. A forged cache_info.n_segs below seg_num otherwise leaves segments[cache_info.n_segs..seg_num) as zeroed structs whose data pointer is NULL, so a forged id in that window would still be dereferenced. A later patch guarantees cache_info.n_segs <= seg_num, and a driver-created cache sets the two equal, so valid images are unaffected. | ||||
| CVE-2026-93179 | 1 Linux | 1 Linux Kernel | 2026-10-15 | N/A |
| In the Linux kernel, the following vulnerability has been resolved: drm/amd/powerplay: fix VoltageObjectInfo zero-stride loop and OOB read Reject voltage objects whose usSize is smaller than the header or would advance the cursor past the table end, preventing an infinite loop or heap OOB read when the VBIOS supplies a malformed VoltageObjectInfo table. | ||||
| CVE-2026-93176 | 1 Linux | 1 Linux Kernel | 2026-10-15 | 7 High |
| In the Linux kernel, the following vulnerability has been resolved: drm/amd/display: Fix dangling pointer in plane reset function amdgpu_dm_plane_drm_plane_reset() frees the old state before allocating a new one. If kzalloc() fails, the function returns without updating the state pointer, leaving a dangling pointer to already freed memory. Fix this by allocating the new state first. On allocation failure, the old state remains untouched and the function safely returns. Found by Linux Verification Center (linuxtesting.org) with SVACE. [adjust for movement around current amd-staging-drm-next] | ||||
| CVE-2026-93175 | 1 Linux | 1 Linux Kernel | 2026-10-15 | 7.8 High |
| In the Linux kernel, the following vulnerability has been resolved: drm/amd/display: Fix dangling pointer in CRTC reset function amdgpu_dm_crtc_reset_state() frees the old state before allocating a new one. If kzalloc() fails, the function returns without updating the state pointer, leaving a dangling pointer to already freed memory. Fix this by allocating the new state first. On allocation failure, the old state remains untouched and the function safely returns. Found by Linux Verification Center (linuxtesting.org) with SVACE. [adjust for movement around current amd-staging-drm-next] | ||||
| CVE-2026-93168 | 1 Linux | 1 Linux Kernel | 2026-10-15 | N/A |
| In the Linux kernel, the following vulnerability has been resolved: dmaengine: xilinx_dma: Fix CPU stall in xilinx_dma_poll_timeout Currently when calling xilinx_dma_poll_timeout with delay_us=0 and a condition that is never fulfilled, the CPU busy-waits for prolonged time and the timeout triggers only with a massive delay causing a CPU stall. This happens due to a huge underestimation of wall clock time in poll_timeout_us_atomic. Commit 7349a69cf312 ("iopoll: Do not use timekeeping in read_poll_timeout_atomic()") changed the behavior to no longer use ktime_get at the expense of underestimation of wall clock time which appears to be very large for delay_us=0. Instead of timing out after approximately XILINX_DMA_LOOP_COUNT microseconds, the timeout takes XILINX_DMA_LOOP_COUNT * 1000 * (time that the overhead of the for loop in poll_timeout_us_atomic takes) which is in the range of several minutes for XILINX_DMA_LOOP_COUNT=1000000. Fix this by using a non-zero value for delay_us. Use delay_us=10 to keep the delay in the hot path of starting DMA transfers minimal but still avoid CPU stalls in case of unexpected hardware failures. One-off measurement with delay_us=0 causes the cpu to busy wait around 7 minutes in the timeout case. After applying this patch with delay_us=10 the measured timeout was 1053428 microseconds which is roughly equivalent to the expected 1000000 microseconds specified in XILINX_DMA_LOOP_COUNT. Add a constant XILINX_DMA_POLL_DELAY_US for delay_us value. | ||||
| CVE-2026-93164 | 1 Linux | 1 Linux Kernel | 2026-10-15 | N/A |
| In the Linux kernel, the following vulnerability has been resolved: uprobes/x86: Move optimized uprobe from nop5 to nop10 Andrii reported an issue with optimized uprobes [1] that can clobber redzone area with call instruction storing return address on stack where user code may keep temporary data without adjusting rsp. Fixing this by moving the optimized uprobes on top of 10-bytes nop instruction, so we can squeeze another instruction to escape the redzone area before doing the call, like: lea -0x80(%rsp), %rsp call tramp Note the lea instruction is used to adjust the rsp register without changing the flags. We use nop10 and following transformation to optimized instructions above and back as suggested by Peterz [2]. Optimize path (int3_update_optimize): 1) Initial state after set_swbp() installed the uprobe: cc 2e 0f 1f 84 00 00 00 00 00 From offset 0 this is INT3 followed by the tail of the original 10-byte NOP. After a previous unoptimization bytes 5..9 may still contain the old call instruction, which remains valid for threads already there. 2) Rewrite the LEA tail and call displacement: cc [8d 64 24 80 e8 d0 d1 d2 d3] From offset 0 this traps on the uprobe INT3. Bytes 1..9 are not executable entry points while byte 0 is trapped. 3) Publish the first LEA byte: [48] 8d 64 24 80 e8 d0 d1 d2 d3 From offset 0 this is: lea -0x80(%rsp), %rsp call <uprobe-trampoline> Unoptimize path (int3_update_unoptimize): 1) Initial optimized state: 48 8d 64 24 80 e8 d0 d1 d2 d3 Same as 3) above. 2) Trap new entries before restoring the NOP bytes: [cc] 8d 64 24 80 e8 d0 d1 d2 d3 From offset 0 this traps. A thread that had already executed the LEA can still reach the intact CALL at offset 5. 3) Restore bytes 1..4 of the original NOP while keeping byte 0 trapped and byte 5 as CALL. cc [2e 0f 1f 84] e8 d0 d1 d2 d3 From offset 0 this still traps. Offset 5 is still the CALL for any thread that was already past the first LEA byte. 4) Publish the first byte of the original NOP: [66] 2e 0f 1f 84 e8 d0 d1 d2 d3 From offset 0 this is the restored 10-byte NOP; the CALL opcode and displacement are now only NOP operands. Offset 5 still decodes as CALL for a thread that was already there. Tthere is only a single target uprobe-trampoline for the given nop10 instruction address, so the CALL instruction will not be changed across unoptimization/optimization cycles. Therefore, any task that is preempted at the CALL instruction is guaranteed to observe that CALL and not anything else. Note as explained in [2] we need to use following nop10: PF1 PF2 ESC NOPL MOD SIB DISP32 NOP10: 0x66, 0x2e, 0x0f, 0x1f, 0x84, 0x00, 0x00, 0x00, 0x00, 0x00 -- cs nopw 0x00000000(%rax,%rax,1) which means we need to allow 0x2e prefix which maps to INAT_PFX_CS attribute in is_prefix_bad function. Also changing the uprobe syscall error when called out of uprobe trampoline to -EPROTO, so we are able to detect the fixed kernel. The optimized uprobe performance stays the same: uprobe-nop : 3.129 ± 0.013M/s uprobe-push : 3.045 ± 0.006M/s uprobe-ret : 1.095 ± 0.004M/s --> uprobe-nop10 : 7.170 ± 0.020M/s uretprobe-nop : 2.143 ± 0.021M/s uretprobe-push : 2.090 ± 0.000M/s uretprobe-ret : 0.942 ± 0.000M/s --> uretprobe-nop10: 3.381 ± 0.003M/s usdt-nop : 3.245 ± 0.004M/s --> usdt-nop10 : 7.256 ± 0.023M/s [1] https://lore.kernel.org/bpf/20260509003146.976844-1-andrii@kernel.org/ [2] https://lore.kernel.org/bpf/20260518104306.GU3102624@noisy.programming.kicks-ass.net/#t | ||||
| CVE-2026-93148 | 1 Linux | 1 Linux Kernel | 2026-10-15 | 7.8 High |
| In the Linux kernel, the following vulnerability has been resolved: bpf: Reject MEM_ALLOC BTF accesses past object bounds BTF struct walks relax the struct-size check for accesses through a trailing flexible array. That is valid for ordinary BTF type walking, but PTR_TO_BTF_ID | MEM_ALLOC values point to objects allocated with the static BTF type size. When walking a MEM_ALLOC object, reject the access before applying the flexible-array relaxation if the access range extends past the struct size. Apply the same policy to struct ID matching so kfunc and kptr type checks do not walk past the allocated object bounds either. | ||||
| CVE-2026-93144 | 1 Linux | 1 Linux Kernel | 2026-10-15 | 7.8 High |
| In the Linux kernel, the following vulnerability has been resolved: bpf: Reject writes through untrusted BTF pointers check_ptr_to_btf_access() lets program-type btf_struct_access callbacks validate writes before the default BTF access path rejects non-read accesses. That bypasses the read-only policy for untrusted BTF pointers created by helpers such as bpf_rdonly_cast(). Reject non-read accesses through PTR_UNTRUSTED BTF pointers at the common entry point, before the callback branch to handle all cases. | ||||
| CVE-2026-93129 | 1 Linux | 1 Linux Kernel | 2026-10-15 | N/A |
| In the Linux kernel, the following vulnerability has been resolved: platform/x86: dell-wmi-base: Fix handling of ultra performance key The commit message of commit 5fbd827eb9c2 ("platform/x86: dell-wmi: Recognise or support new switches") states that the ultra performance key contains additional data after the type and code fields. The event data passed to dell_wmi_process_key() is already parsed, so "buffer" already starts after those two fields. Use the correct index for accessing the first data field to avoid a potential buffer overread. | ||||
| CVE-2026-93125 | 1 Linux | 1 Linux Kernel | 2026-10-15 | 7.8 High |
| In the Linux kernel, the following vulnerability has been resolved: bpf: Reject rdonly/rdwr_buf_size kfunc arguments that exceed u32 max check_kfunc_args() detects a kfunc argument named rdonly_buf_size or rdwr_buf_size and stores reg->var_off.value into meta->r0_size, a u64, and does not bound it. check_kfunc_call() later copies that value into the returned register's mem_size field: meta->r0_size = reg->var_off.value; ... regs[BPF_REG_0].mem_size = meta.r0_size; regs[BPF_REG_0].mem_size is u32. A constant whose upper 32 bits are set gets truncated instead of causing a load-time rejection, so the verifier records a PTR_TO_MEM register with an approximately 4 GiB mem_size for whatever allocation the kfunc returned. A later access check against that register uses the truncated, wrong bound. Reject rdonly_buf_size/rdwr_buf_size values that exceed U32_MAX at the point meta->r0_size is set. | ||||
| CVE-2026-93122 | 1 Linux | 1 Linux Kernel | 2026-10-15 | 7.8 High |
| In the Linux kernel, the following vulnerability has been resolved: usb: gadget: uac: validate rate list length before storing UAC1 and UAC2 configfs rate-list attributes parse a comma-separated list of sampling rates and store each parsed value in fixed-size arrays. The arrays have UAC_MAX_RATES entries, but the store paths do not check that the input contains at most that many tokens before writing through opts->name##s[i++]. Writing more than ten rates therefore writes past the end of the p_srates[] or c_srates[] array in struct f_uac1_opts or struct f_uac2_opts. With CONFIG_UBSAN_BOUNDS enabled, writing an 11-entry rate list to the UAC1 p_srate attribute reports: UBSAN: array-index-out-of-bounds drivers/usb/gadget/function/f_uac1.c:1669:1 index 10 is out of range for type 'int [10]' __ubsan_handle_out_of_bounds.cold f_uac1_opts_p_srate_store configfs_write_iter vfs_write ksys_write do_syscall_64 The same reproducer against the UAC2 p_srate attribute reports: UBSAN: array-index-out-of-bounds drivers/usb/gadget/function/f_uac2.c:2087:1 index 10 is out of range for type 'int [10]' __ubsan_handle_out_of_bounds.cold f_uac2_opts_p_srate_store configfs_write_iter vfs_write ksys_write do_syscall_64 Reject additional tokens once UAC_MAX_RATES entries have been parsed. Also keep the original kstrdup() pointer for kfree(), because strsep() advances the parsing cursor. Freeing the advanced cursor leaks the original buffer on successful parses and can free an interior pointer on some error paths. | ||||
| CVE-2026-93116 | 1 Linux | 1 Linux Kernel | 2026-10-15 | 7 High |
| In the Linux kernel, the following vulnerability has been resolved: platform/x86: asus-wmi: fix resource leaks on probe failure During driver initialization in asus_wmi_add(), various subsystems are registered sequentially. However, the error path labels are out of order relative to the registration sequence. Specifically: 1. If asus_wmi_custom_fan_curve_init() fails, the driver jumps to fail_custom_fan_curve. Because this label is placed below fail_sysfs, it bypasses the cleanup calls for the input device and sysfs groups, which were successfully registered before, leaking those resources. 2. If asus_screenpad_init() fails, the driver jumps to fail_screenpad. Because fail_screenpad is placed below fail_backlight, it bypasses the cleanup calls for backlight and rfkill, leaking those resources. Fix these resource leaks by reordering the error path labels in asus_wmi_add() to match the exact reverse order of the resource allocations. | ||||
| CVE-2026-93105 | 1 Linux | 1 Linux Kernel | 2026-10-15 | 7.8 High |
| In the Linux kernel, the following vulnerability has been resolved: esp: do not unref managed frag pages in esp_ssg_unref() esp_ssg_unref() releases the page references held on the source scatterlist after the AEAD operation completes. It calls skb_page_unref() on every frag page for an out-of-place transform (req->src != req->dst), and in the error path of esp_output_tail() (already_unref == true) on the request's own scatterlist. This is wrong when the skb carries managed frags (SKBFL_MANAGED_FRAG_REFS). Managed frags are owned by a zerocopy ubuf and the skb does not hold a per-frag page reference; io_uring SEND_ZC with a registered buffer attaches the bvec pages this way via io_sg_from_iter(). The rest of the stack honours this invariant: skb_release_data() skips the per-frag unref when SKBFL_MANAGED_FRAG_REFS is set, and skb_zcopy_managed() is the guard used at the other unref sites. esp_ssg_unref() is missing that guard, so for a managed-frag skb it drops a page reference the skb never acquired. This can underflow the page reference count and free a page that is still in use. Guard the function with skb_zcopy_managed() so both unref paths are skipped for managed-frag skbs, matching skb_release_data(). | ||||
| CVE-2026-93104 | 1 Linux | 1 Linux Kernel | 2026-10-15 | N/A |
| In the Linux kernel, the following vulnerability has been resolved: RDMA/rvt: Return NULL after port allocation failure rvt_alloc_device() deallocates the IB device when its port array cannot be allocated but then returns the pointer to the released allocation. Callers treat any non-NULL value as valid and dereference it, resulting in a use-after-free. Return NULL immediately after deallocation so callers can propagate the allocation failure. | ||||
| CVE-2026-93099 | 1 Linux | 1 Linux Kernel | 2026-10-15 | N/A |
| In the Linux kernel, the following vulnerability has been resolved: fs/resctrl: Fix UAF from worker threads when domains are removed The mbm_handle_overflow() and cqm_handle_limbo() workers read event counters and may sleep while doing so. They are scheduled via delayed_work embedded in struct rdt_l3_mon_domain. Architecture allocates and frees these domains from CPU hotplug callbacks under cpus_write_lock(), and the workers acquire cpus_read_lock() to keep the domain alive across their access. A use-after-free can occur when a worker is blocked waiting for cpus_read_lock() while the hotplug core holds cpus_write_lock(): the architecture frees the rdt_l3_mon_domain that contains the worker's work_struct. When the worker unblocks, the container_of() it performs on the embedded work pointer dereferences freed memory. Drop cpus_read_lock() from the workers and instead drain pending and in-flight work synchronously before the architecture can free the domain. Since architecture offlines the domain under cpus_write_lock() after it has been unlinked from the RCU list and a grace period has elapsed, no new work can be scheduled. The cancel only needs to wait out existing work. Drop rdtgroup_mutex during CPU offline around cancel_delayed_work_sync() so that a worker waiting on the mutex can complete before re-pinning the work on a different CPU. When offlining a CPU the architecture may iterate over resources in any order. For example, the MBA control domain may be offlined before or after a corresponding L3 monitor domain. Ensure that resctrl fs cancels the workers no matter what order the architecture offlines the domains. | ||||
| CVE-2026-93096 | 1 Linux | 1 Linux Kernel | 2026-10-15 | N/A |
| In the Linux kernel, the following vulnerability has been resolved: cxl/features: Serialize multi-part Get/Set Feature transfers A Get or Set Feature payload larger than the mailbox payload size is split into several mailbox commands. mbox_mutex only serializes individual mailbox commands and is dropped between iterations of these loops. Nothing serializes the multi-part transfer as a whole. cxl_get_feature() and cxl_set_feature() are reachable concurrently from fwctl (per-fd RPCs run under a read-held registration lock) and from the EDAC scrub/ECS/repair paths, so two transfers to the same mailbox can interleave their parts and corrupt the device's transfer context. Add a per-mailbox feat_mutex and hold it across the whole transfer in both functions. It nests outside mbox_mutex (which is taken inside cxl_internal_send_cmd()), and is taken nowhere else, so no lock-ordering inversion is introduced. | ||||