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| CVE | Vendors | Products | Updated | CVSS v3.1 |
|---|---|---|---|---|
| CVE-2026-47870 | 2026-07-18 | 7.1 High | ||
| VMware Avi Load Balancer contains a privilege escalation vulnerability. A malicious authenticated user with network access may be able to execute remote code. Affected versions: 32.1.1 (fixed in 32.1.2) 31.1.1 through 31.2.2 (fixed in 31.2.2-2p3) 30.1.1 through 30.2.6 (fixed in 30.2.7) 22.1.1 through 22.1.7 (fixed in 30.2.7) | ||||
| CVE-2026-47869 | 2026-07-18 | 8.7 High | ||
| VMware Avi Load Balancer contains a remote code execution vulnerability. A malicious authenticated user with network access may be able to inject and execute code. Affected versions: 32.1.1 (fixed in 32.1.2) 31.1.1 through 31.2.2 (fixed in 31.2.2-2p3) 30.1.1 through 30.2.6 (fixed in 30.2.7) 22.1.1 through 22.1.7 (fixed in 30.2.7) | ||||
| CVE-2026-47868 | 2026-07-18 | 7.8 High | ||
| VMware Avi Load Balancer contains a local privilege escalation vulnerability. A malicious user with local access may be able to escalate their privileges to run code as root. Affected versions: 32.1.1 (fixed in 32.1.2) 31.1.1 through 31.2.2 (fixed in 31.2.2-2p3) 30.1.1 through 30.2.6 (fixed in 30.2.7) 22.1.1 through 22.1.7 (fixed in 30.2.7) | ||||
| CVE-2026-47867 | 2026-07-18 | 8.7 High | ||
| VMware Avi Load Balancer contains a remote code execution vulnerability. A malicious user with network access may be able to access the Avi Control plane and execute code remotely. Affected versions: 32.1.1 (fixed in 32.1.2) 31.1.1 through 31.2.2 (fixed in 31.2.2-2p3) 30.1.1 through 30.2.6 (fixed in 30.2.7) 22.1.1 through 22.1.7 (fixed in 30.2.7) | ||||
| CVE-2026-47866 | 2026-07-18 | 8.3 High | ||
| VMware Avi Load Balancer contains an authorization bypass vulnerability. A malicious actor on the network can access a limited subset of the Avi Control Plane without proper authorization. Affected versions: 32.1.1 (fixed in 32.1.2) 31.1.1 through 31.2.2 (fixed in 31.2.2-2p3) 30.1.1 through 30.2.6 (fixed in 30.2.7) 22.1.1 through 22.1.7 (fixed in 30.2.7) | ||||
| CVE-2026-47865 | 2026-07-18 | 9.8 Critical | ||
| VMware Avi Load Balancer contains an authentication bypass vulnerability. A malicious user with network access may be able to access the Avi Control plane by bypassing the authentication mechanism. Affected versions: 31.1.1 through 31.2.2 (fixed in 31.2.2-2p3) 30.1.1 through 30.2.6 (fixed in 30.2.7) 22.1.1 through 22.1.7 (fixed in 30.2.7) | ||||
| CVE-2026-16083 | 1 Sipeed | 1 Picoclaw | 2026-07-18 | 5.3 Medium |
| A security flaw has been discovered in Sipeed PicoClaw up to 0.2.9. This affects the function webhook.ParseRequest of the file pkg/channels/line/line.go of the component LINE Webhook. The manipulation results in authentication bypass by capture-replay. The attack may be launched remotely. The exploit has been released to the public and may be used for attacks. The reported GitHub issue was closed automatically with the label "not planned" by a bot. | ||||
| CVE-2026-16082 | 1 Sipeed | 1 Picoclaw | 2026-07-18 | 5.3 Medium |
| A vulnerability was identified in Sipeed PicoClaw up to 0.2.9. The impacted element is the function ExecTool.executeRun of the file pkg/agent/pipeline_execute.go. The manipulation of the argument cwe leads to time-of-check time-of-use. The attack must be carried out locally. The exploit is publicly available and might be used. The reported GitHub issue was closed automatically with the label "not planned" by a bot. | ||||
| CVE-2026-16081 | 1 Sipeed | 1 Picoclaw | 2026-07-18 | 4.3 Medium |
| A vulnerability was determined in Sipeed PicoClaw up to 0.2.9. The affected element is an unknown function of the file web/backend/api/auth.go. Executing a manipulation can lead to cross-site request forgery. The attack can be launched remotely. The exploit has been publicly disclosed and may be utilized. This patch is called 4b0229351678f479429b8d8b19207757266f246b. Applying a patch is advised to resolve this issue. | ||||
| CVE-2026-53366 | 2 Linux, Redhat | 3 Linux Kernel, Enterprise Linux, Enterprise Linux Eus | 2026-07-18 | 7.8 High |
| In the Linux kernel, the following vulnerability has been resolved: ipv4: account for fraggap on the paged allocation path In __ip_append_data(), when the paged-allocation branch is taken, alloclen and pagedlen are computed as alloclen = fragheaderlen + transhdrlen; pagedlen = datalen - transhdrlen; datalen already includes fraggap, but the fraggap bytes carried over from the previous skb are copied into the new skb's linear area at offset transhdrlen by the subsequent skb_copy_and_csum_bits(). The linear area is therefore undersized by fraggap bytes while pagedlen is overstated by the same amount. The non-paged branch sets alloclen to fraglen, which already accounts for fraggap because datalen does. Bring the paged branch in line by adding fraggap to alloclen and subtracting it from pagedlen. After this adjustment, copy no longer collapses to -fraggap on the paged path, so remove the stale comment describing that old arithmetic. | ||||
| CVE-2026-53363 | 1 Linux | 1 Linux Kernel | 2026-07-18 | 9.8 Critical |
| In the Linux kernel, the following vulnerability has been resolved: xfrm: iptfs: preserve shared-frag marker in iptfs_consume_frags() iptfs_consume_frags() transfers paged fragments from one socket buffer to another but fails to propagate the SKBFL_SHARED_FRAG flag. This is the same class of bug that was fixed in skb_try_coalesce() for CVE-2026-46300: when fragments backed by read-only page-cache pages are merged, the marker indicating their shared nature must be preserved so that ESP can decide correctly whether in-place encryption is safe. Apply the same two-line fix used in skb_try_coalesce() to iptfs_consume_frags(). | ||||
| CVE-2026-53362 | 1 Linux | 1 Linux Kernel | 2026-07-18 | 7.8 High |
| In the Linux kernel, the following vulnerability has been resolved: ipv6: account for fraggap on the paged allocation path In __ip6_append_data(), when the paged-allocation branch is taken (MSG_MORE / NETIF_F_SG / large fraglen), alloclen and pagedlen are computed as alloclen = fragheaderlen + transhdrlen; pagedlen = datalen - transhdrlen; datalen already includes fraggap (datalen = length + fraggap). When fraggap is non-zero, this is not the first skb and transhdrlen is zero. The fraggap bytes carried over from the previous skb are copied just past the fragment headers in the new skb's linear area. The linear area is therefore undersized by fraggap bytes while pagedlen is overstated by the same amount, and the copy writes past skb->end into the trailing skb_shared_info. An unprivileged user can trigger this via a UDPv6 socket using MSG_MORE together with MSG_SPLICE_PAGES. The bad accounting was introduced by commit 773ba4fe9104 ("ipv6: avoid partial copy for zc"). Before commit ce650a166335 ("udp6: Fix __ip6_append_data()'s handling of MSG_SPLICE_PAGES"), the negative copy value caused -EINVAL to be returned. That later commit allowed MSG_SPLICE_PAGES to proceed in this case, making the corruption triggerable. The non-paged branch sets alloclen to fraglen, which already accounts for fraggap because datalen does. Bring the paged branch in line by adding fraggap to alloclen and subtracting it from pagedlen. After this adjustment, copy no longer collapses to -fraggap on the paged path, so remove the stale comment describing that old arithmetic. Since a negative copy is no longer expected for a valid MSG_SPLICE_PAGES case, remove the MSG_SPLICE_PAGES exception from the negative copy check. | ||||
| CVE-2026-53361 | 1 Linux | 1 Linux Kernel | 2026-07-18 | 7.1 High |
| In the Linux kernel, the following vulnerability has been resolved: af_unix: Set gc_in_progress to true in unix_gc(). Igor Ushakov reported that unix_gc() could run with gc_in_progress being false if the work is scheduled while running: Thread 1 Thread 2 Thread 3 -------- -------- -------- unix_schedule_gc() unix_schedule_gc() `- if (!gc_in_progress) `- if (!gc_in_progress) |- gc_in_progress = true | `- queue_work() | unix_gc() <----------------/ | | |- gc_in_progress = true ... `- queue_work() | | `- gc_in_progress = false | | unix_gc() <---------------------------------------------' | ... /* gc_in_progress == false */ | `- gc_in_progress = false unix_peek_fpl() relies on gc_in_progress not to confuse GC by MSG_PEEK. Let's set gc_in_progress to true in unix_gc(). | ||||
| CVE-2026-53360 | 1 Linux | 1 Linux Kernel | 2026-07-18 | 8.8 High |
| In the Linux kernel, the following vulnerability has been resolved: KVM: SEV: Require in-GHCB scratch area if GHCB v2+ is in use As per the GHCB spec, when using GHCB v2+ require the software scratch area to reside in the GHCB's shared buffer. Note, things like Page State Change (PSC) requests _rely_ on this behavior, as the guest can't provide a length when making the request, i.e. the size of the guest payload is bounded by the size of the shared buffer. Failure to force usage of the GHCB, and a slew of other flaws, lets a malicious SNP guest corrupt host kernel heap memory, and leak host heap layout information. setup_vmgexit_scratch() allocates a buffer via kvzalloc(exit_info_2), where exit_info_2 is guest-controlled. With exit_info_2=24, this yields a 24-byte allocation in kmalloc-cg-32 (32-byte slab objects). The buffer holds an 8-byte psc_hdr followed by 8-byte psc_entry structs, so only entries[0] and entries[1] are in-bounds. snp_begin_psc() validates end_entry against VMGEXIT_PSC_MAX_COUNT (253) but NOT against the actual buffer size: idx_end = hdr->end_entry; if (idx_end >= VMGEXIT_PSC_MAX_COUNT) { // checks 253, not buffer snp_complete_psc(svm, ...); return 1; } for (idx = idx_start; idx <= idx_end; idx++) { entry_start = entries[idx]; // OOB when idx >= 2 The guest sets end_entry=10+, causing the host to iterate entries[2+] which are OOB into adjacent slab objects. For each OOB entry: - The host reads 8 bytes (OOB READ / info leak oracle) - If the data passes PSC validation, __snp_complete_one_psc() writes cur_page = 1 or 512 into the entry (OOB WRITE, sev.c:3806) - If validation fails, the error response reveals whether adjacent memory is zero vs non-zero (information disclosure to guest) The guest controls allocation size (exit_info_2), entry range (cur_entry/end_entry), and can fire unlimited VMGEXITs to repeatedly hit different slab positions. By exploiting the variety of bugs, a malicious SEV-SNP guest can: - OOB read adjacent kmalloc-cg-32 objects (heap layout disclosure) - OOB write cur_page bits into adjacent objects (heap corruption) - Trigger use-after-free conditions across VMGEXITs E.g. with KASAN enabled, a single insmod of the PoC guest module produces 73 KASAN reports: BUG: KASAN: slab-out-of-bounds in snp_begin_psc+0x126/0x890 Read of size 8 at addr ffff888219ffb5e0 by task qemu-system-x86/2199 BUG: KASAN: slab-out-of-bounds in snp_begin_psc+0x468/0x890 Write of size 8 at addr ffff888351566648 by task qemu-system-x86/2199 The buggy address belongs to the object at ffff888XXXXXXXXX which belongs to the cache kmalloc-cg-32 of size 32 The buggy address is located N bytes to the right of allocated 32-byte region [ffff888XXXXXXXXX, ffff888XXXXXXXXX) Breakdown: 62 slab-out-of-bounds (reads + writes past allocation) 7 slab-use-after-free 4 use-after-free All credit to Stan for the wonderful description and reproducer! [sean: write changelog] | ||||
| CVE-2026-53359 | 1 Linux | 1 Linux Kernel | 2026-07-18 | 8.8 High |
| In the Linux kernel, the following vulnerability has been resolved: KVM: x86: Fix shadow paging use-after-free due to unexpected role Commit 0cb2af2ea66ad ("KVM: x86: Fix shadow paging use-after-free due to unexpected GFN") fixed a shadow paging mismatch between stored and computed GFNs; the bug could be triggered by changing a PDE mapping from outside the guest, and then deleting a memslot. The rmap_remove() call would miss entries created after the PDE change because the GFN of the leaf SPTE does not match the GFN of the struct kvm_mmu_page. A similar hole however remains if the modified PDE points to a non-leaf page. In this case the gfn can be made to match, but the role does not match: the original large 2MB page creates a kvm_mmu_page with direct=1, while the new 4KB needs a kvm_mmu_page with direct=0. However, kvm_mmu_get_child_sp() does not compare the role, and therefore reuses the page. The next step is installing a leaf (4KB) SPTE on the new path which records an rmap entry under the gfn resolved by the walk. But when that child is zapped its parent kvm_mmu_page has direct=1 and kvm_mmu_page_get_gfn() computes the gfn for the 4KB page as sp->gfn + index instead of using sp->shadowed_translation[] (or sp->gfns[] in older kernels). It therefore fails to remove the recorded entry. When the memslot is dropped the shadow page is freed but the rmap entry survives, as in the scenario that was already fixed. Code that later walks that gfn (dirty logging, MMU notifier invalidation, and so on) dereferences an sptep that lies in the freed page, causing the use-after-free. | ||||
| CVE-2026-53358 | 1 Linux | 1 Linux Kernel | 2026-07-18 | 8.8 High |
| In the Linux kernel, the following vulnerability has been resolved: Bluetooth: L2CAP: use chan timer to close channels in cleanup_listen() l2cap_chan_close() removes the channel from conn->chan_l, which must be done under conn->lock. cleanup_listen() runs under the parent sk_lock, so acquiring conn->lock would invert the established conn->lock -> chan->lock -> sk_lock order. Instead of calling l2cap_chan_close() directly, schedule l2cap_chan_timeout with delay 0 to close the channel asynchronously. The timeout handler already acquires conn->lock and chan->lock in the correct order. The timer is only armed when chan->conn is still set: if it is already NULL, l2cap_conn_del() has already processed this channel (l2cap_chan_del + l2cap_sock_teardown_cb + l2cap_sock_close_cb), so there is nothing left to do. If l2cap_conn_del() races in after the timer is armed, __clear_chan_timer() inside l2cap_chan_del() cancels it; if the timer has already fired, the handler returns harmlessly because chan->conn was cleared. | ||||
| CVE-2026-53357 | 1 Linux | 1 Linux Kernel | 2026-07-18 | 8 High |
| In the Linux kernel, the following vulnerability has been resolved: Bluetooth: fix UAF in l2cap_sock_cleanup_listen() vs l2cap_conn_del() bt_accept_dequeue() unlinks a not-yet-accepted child from the parent accept queue and release_sock()s it before returning, so the returned sk has no caller reference and is unlocked. l2cap_sock_cleanup_listen() walks these children on listening-socket close. A concurrent HCI disconnect drives hci_rx_work -> l2cap_conn_del() which runs l2cap_chan_del() + l2cap_sock_kill() and frees the child sk and its l2cap_chan; cleanup_listen() then uses both: BUG: KASAN: slab-use-after-free in l2cap_sock_kill l2cap_sock_kill / l2cap_sock_cleanup_listen / __x64_sys_close Freed by: l2cap_conn_del -> l2cap_sock_close_cb -> l2cap_sock_kill This is distinct from the two fixes already in this area: commit e83f5e24da741 ("Bluetooth: serialize accept_q access") serialises the accept_q list/poll and takes temporary refs inside bt_accept_dequeue(), and CVE-2025-39860 serialises the userspace close()/accept() race by calling cleanup_listen() under lock_sock() in l2cap_sock_release(). Neither covers l2cap_conn_del() running from hci_rx_work, so this UAF still reproduces on current bluetooth/master. Take the reference at the source: bt_accept_dequeue() does sock_hold() while sk is still locked, before release_sock(); callers sock_put(). cleanup_listen() pins the chan with l2cap_chan_hold_unless_zero() under a brief child sk lock (serialising vs l2cap_sock_teardown_cb()), drops it before l2cap_chan_lock(), and skips a duplicate l2cap_sock_kill() on SOCK_DEAD. conn->lock is not taken here: cleanup_listen() runs under the parent sk lock and that would invert conn->lock -> chan->lock -> sk_lock (lockdep). KASAN/SMP: an unprivileged listen/close vs HCI-disconnect race produced 12 use-after-free reports per run before this change; 0, and no lockdep report, over 1600+ raced iterations after it on bluetooth/master. | ||||
| CVE-2026-53356 | 1 Linux | 1 Linux Kernel | 2026-07-18 | 7.8 High |
| In the Linux kernel, the following vulnerability has been resolved: drm/i915/gem: Fix phys BO pread/pwrite with offset sg_page() returns struct page pointer not (void *) so the scaling of pread/pwrite is wrong for phys BO and wrong parts of BO would be accessed if non-zero offset is used. Last impacted platform with overlay or cursor planes using phys mapping was Gen3/945G/Lakeport. (cherry picked from commit 3e49a2f85070b2fb672c1e0fdba281a4ea3aebe6) | ||||
| CVE-2026-53355 | 1 Linux | 1 Linux Kernel | 2026-07-18 | 9.8 Critical |
| In the Linux kernel, the following vulnerability has been resolved: net: rds: clear i_sends on setup unwind The RDS IB connection teardown path is written so it can run during partial startup and on repeated shutdown attempts. It uses NULL pointers to distinguish resources that are still owned from resources that have already been released. When rds_ib_setup_qp() fails after allocating i_sends but before allocating i_recvs, the sends_out path frees i_sends without clearing the pointer. A later shutdown pass can still treat that stale pointer as a live send ring allocation. Clear i_sends after vfree() in the error unwind path so the existing shutdown logic continues to use the correct ownership state. | ||||
| CVE-2026-53354 | 1 Linux | 1 Linux Kernel | 2026-07-18 | 8.8 High |
| In the Linux kernel, the following vulnerability has been resolved: arm64: errata: Mitigate TLBI errata on various Arm CPUs A number of CPUs developed by Arm suffer from errata whereby a broadcast TLBI;DSB sequence may complete before the global observation of writes which are translated by an affected TLB entry. These errata ONLY affect the completion of memory accesses which have been translated by an invalidated TLB entry, and these errata DO NOT affect the actual invalidation of TLB entries. TLB entries are removed correctly. This issue has been assigned CVE ID CVE-2025-10263. To mitigate this issue, Arm recommends that software follows any affected TLBI;DSB sequence with an additional TLBI;DSB, which will ensure that all memory write effects affected by the first TLBI have been globally observed. The additional TLBI can use any operation that is broadcast to affected CPUs, and the additional DSB can use any option that is sufficient to complete the additional TLBI. The ARM64_WORKAROUND_REPEAT_TLBI workaround is sufficient to mitigate the issue. Enable this workaround for affected CPUs, and update the silicon errata documentation accordingly. Note that due to the manner in which Arm develops IP and tracks errata, some CPUs share a common erratum number. | ||||