CVE Database

In the Linux kernel, the following vulnerability has been resolved: isofs: bound Rock Ridge symlink components to the SL record get_symlink_chunk() and the SL handling in parse_rock_ridge_inode_internal() walk the variable-length components of a Rock Ridge "SL" (symbolic link) record. Each component is a two-byte header (flags, len) followed by len bytes of text, so it occupies slp->len + 2 bytes. Both loops read slp->len and advance to the next component, and get_symlink_chunk() additionally does memcpy(rpnt, slp->text, slp->len), but neither checks that the component lies within the SL record before dereferencing it. A crafted SL record whose component declares a len that runs past the record (rr->len) therefore triggers an out-of-bounds read of up to 255 bytes. When the record sits at the tail of its backing buffer - for example a small kmalloc()ed continuation block reached through a CE record - the read crosses the allocation; get_symlink_chunk() then copies the out-of-bounds bytes into the symlink body returned to user space by readlink(), disclosing adjacent kernel memory. ISO 9660 images are routinely mounted from untrusted removable media - desktop environments auto-mount them (e.g. via udisks2) without CAP_SYS_ADMIN - so the record contents are attacker-controlled. Reject any component that does not fit in the remaining record bytes before using it. In get_symlink_chunk() return NULL, like the existing output-buffer (plimit) checks, so a malformed record makes readlink() fail with -EIO rather than silently returning a truncated target; in parse_rock_ridge_inode_internal() stop the inode-size walk.

In the Linux kernel, the following vulnerability has been resolved: crypto: caam - use print_hex_dump_devel to guard key hex dumps Use print_hex_dump_devel() for dumping sensitive key material in *_setkey() and gen_split_key() to avoid leaking secrets at runtime when CONFIG_DYNAMIC_DEBUG is enabled.

In the Linux kernel, the following vulnerability has been resolved: crypto: caam - use print_hex_dump_devel to guard key hex dumps Use print_hex_dump_devel() for dumping sensitive key material in *_setkey() to avoid leaking secrets at runtime when CONFIG_DYNAMIC_DEBUG is enabled.

In the Linux kernel, the following vulnerability has been resolved: crypto: chacha20poly1305 - validate poly1305 template argument chachapoly_create() still accepts the compatibility poly1305 parameter in the template name, but it assumes the second template argument is always present and immediately passes it to strcmp(). When the argument is missing, crypto_attr_alg_name() returns an error pointer. Check for that before comparing the name so malformed template instantiations fail with an error instead of dereferencing the error pointer in strcmp(). This matches the surrounding Crypto API template pattern where crypto_attr_alg_name() results are validated before string-specific use.

In the Linux kernel, the following vulnerability has been resolved: crypto: ecc - Fix carry overflow in vli multiplication The carry flag calculation fails when r01.m_high is saturated (0xFFFFFFFFFFFFFFFF) and addition of lower bits overflows. The condition (r01.m_high < product.m_high) doesn't handle the case where r01.m_high == product.m_high and an additional carry exists from lower-bit overflow. When commit 3c4b23901a0c ("crypto: ecdh - Add ECDH software support") introduced crypto/ecc.c, it split the muladd() function in the micro-ecc library into separate mul_64_64() and add_128_128() helpers. It seems the check got lost in translation. Add proper handling for this boundary by accounting for the carry from the lower addition.

In the Linux kernel, the following vulnerability has been resolved: crypto: pcrypt - restore callback for non-parallel fallback pcrypt installs pcrypt_aead_done() on the child AEAD request before trying to submit it through padata. If padata_do_parallel() returns -EBUSY, pcrypt falls back to calling the child AEAD directly. That fallback must not keep the padata completion callback. Otherwise an asynchronous completion runs pcrypt_aead_done() even though the request was never enrolled in padata. Restore the original request callback and callback data before calling the child AEAD directly. This keeps the fallback path aligned with a direct AEAD request while leaving the parallel path unchanged.

In the Linux kernel, the following vulnerability has been resolved: crypto: loongson - Remove broken and unused loongson-rng The loongson-rng rng_alg has several vulnerabilities, including not providing forward security, and a use-after-free bug due to the use of wait_for_completion_interruptible(). Meanwhile, the rng_alg framework doesn't really have any purpose in the first place other than to access the software algorithms crypto/drbg.c and crypto/jitterentropy.c. Hardware-specific rng_algs have no in-kernel user, and unlike hwrng there's no feed into the actual Linux RNG. As such, there's really no point to this code. There are of course other rng_alg drivers that are similarly unused, but they're similarly in the process of being phased out, e.g. https://lore.kernel.org/r/20260529193648.18172-1-ebiggers@kernel.org and https://lore.kernel.org/r/20260529220430.34135-1-ebiggers@kernel.org Given that, there's no point in fixing forward these vulnerabilities, and it makes much more sense to simply roll back the addition of this driver. If this platform provides TRNG (not PRNG) functionality, it could make sense to add a hwrng driver, but it would be quite different.

In the Linux kernel, the following vulnerability has been resolved: crypto: ccp - Do not initialize SNP for SEV ioctls Sashiko notes: > if SEV initialization fails and KVM is actively running normal VMs, could a > userspace process trigger this code path via /dev/sev ioctls (e.g., > SEV_PDH_GEN) and zero out MSR_VM_HSAVE_PA globally? Would the next VMRUN > execution for an active VM trigger a general protection fault and crash the > host? sev_move_to_init_state() is called for ioctls requiring only SEV firmware: SEV_PEK_GEN, SEV_PDH_GEN, SEV_PEK_CSR, SEV_PEK_CERT_IMPORT, and SEV_PDH_CERT_EXPORT. After the firmware command, it does SEV_SHUTDOWN on the SEV firmware. Since these commands do not require SNP to be initialized, skip it by calling __sev_platform_init_locked() which only initializes the SEV firmware. This way SNP is not Initialized at all, and HSAVE_PA is not cleared. The previous code saved any SEV initialization firmware error to init_args.error and then threw it away and hardcoded the return value of INVALID_PLATFORM_STATE regardless of the real firmware error. This patch changes it to surface the underlying error, which is hopefully both more useful and doesn't cause any problems. Note that it is still safe to call __sev_firmware_shutdown() directly: it calls __sev_snp_shutdown_locked(), which skips SNP shutdown if SNP was not initialized.

In the Linux kernel, the following vulnerability has been resolved: crypto: ccp - Do not initialize SNP for ioctl(SNP_COMMIT) Sashiko notes: > if SEV initialization fails and KVM is actively running normal VMs, could a > userspace process trigger this code path via /dev/sev ioctls (e.g., > SEV_PDH_GEN) and zero out MSR_VM_HSAVE_PA globally? Would the next VMRUN > execution for an active VM trigger a general protection fault and crash the > host? The SNP_COMMIT command does not require the firmware to be in any particular state. Skip initializing it if it was previously uninitialized. The SEV-SNP firmware specification doc 56860 does not mention SNP_COMMIT in Table 5 as a command that is allowed in the UNINIT state, but it is in fact allowed and a future documentation update will reflect that.

In the Linux kernel, the following vulnerability has been resolved: crypto: ccp - Do not initialize SNP for ioctl(SNP_VLEK_LOAD) Sashiko notes: > if SEV initialization fails and KVM is actively running normal VMs, could a > userspace process trigger this code path via /dev/sev ioctls (e.g., > SEV_PDH_GEN) and zero out MSR_VM_HSAVE_PA globally? Would the next VMRUN > execution for an active VM trigger a general protection fault and crash the > host? The SEV firmware docs for SNP_VLEK_LOAD note: > On SNP_SHUTDOWN, the VLEK is deleted. That is, the initialization/shutdown wrapper here is pointless, because the firmware immediately throws away the key anyway. Instead, refuse to do anything if SNP has not been previously initialized. This is an ABI break: before, this was a no-op and almost certainly a mistake by userspace, and now it returns -ENODEV. ABI compatibility could be maintained here by simply returning 0 in the check instead.

In the Linux kernel, the following vulnerability has been resolved: crypto: ccp - Do not initialize SNP for ioctl(SNP_CONFIG) Sashiko notes: > if SEV initialization fails and KVM is actively running normal VMs, could a > userspace process trigger this code path via /dev/sev ioctls (e.g., > SEV_PDH_GEN) and zero out MSR_VM_HSAVE_PA globally? Would the next VMRUN > execution for an active VM trigger a general protection fault and crash the > host? Refuse to re-try initialization if SNP is not already initialized for SNP_CONFIG. This is technically an ABI break: before if SNP initialization failed it could be transparently retriggered by this ioctl, and if no VMs were running, everything worked fine. Hopefully this is enough of a corner case that nobody will notice, but someone does, there are a few options: * do something like symbol_get() for kvm and refuse to initialize if KVM is loaded * check each cpu's HSAVE_PA for non-zero data before re-initializing * once initialization has failed, continue to refuse to initialize until the ccp module is unloaded

In the Linux kernel, the following vulnerability has been resolved: crypto: drbg - Fix returning success on failure in CTR_DRBG drbg_ctr_generate() sometimes returns success when it fails, leaving the output buffer uninitialized. Fix it.

In the Linux kernel, the following vulnerability has been resolved: crypto: qat - protect service table iterations with service_lock The service_table list is protected by service_lock when entries are added or removed (in adf_service_add() and adf_service_remove()), but several functions iterate over the list without holding this lock. A concurrent adf_service_register() or adf_service_unregister() call could modify the list during traversal, leading to list corruption or a use-after-free. Fix this by holding service_lock across all list_for_each_entry() iterations of service_table in adf_dev_init(), adf_dev_start(), adf_dev_stop(), adf_dev_shutdown(), adf_dev_restarting_notify(), adf_dev_restarted_notify(), and adf_error_notifier(). The lock ordering is safe: callers of the static helpers (adf_dev_up() and adf_dev_down()) acquire state_lock before service_lock, and no event_hld callback or service_lock holder ever acquires state_lock in the reverse order.

In the Linux kernel, the following vulnerability has been resolved: crypto: qat - validate RSA CRT component lengths The generic RSA key parser (rsa_helper.c) bounds each CRT component (p, q, dp, dq, qinv) by the modulus size n_sz, but qat_rsa_setkey_crt() allocates half-size DMA buffers (key_sz / 2) and right-aligns each component with: memcpy(dst + half_key_sz - len, src, len) When a CRT component is larger than half_key_sz the subtraction underflows and memcpy writes past the DMA buffer, causing memory corruption. Add a len > half_key_sz check next to the existing !len check for each of the five CRT components so the driver falls back to the non-CRT path instead of writing out of bounds.

In the Linux kernel, the following vulnerability has been resolved: spi: fsl-lpspi: terminate the RX channel on TX prepare failure path When dmaengine_prep_slave_sg() fails for the TX channel, the error path terminates the TX DMA channel but leaves the RX channel running. Since the RX channel was already submitted and issued prior to preparing the TX descriptor, returning -EINVAL causes the SPI core to unmap the DMA buffers while the RX DMA engine continues writing to them, leading to potential memory corruption or use-after-free. Terminate the RX channel before returning on the TX prepare failure path.

In the Linux kernel, the following vulnerability has been resolved: x86/mm: Fix freeing of PMD-sized vmemmap pages Commit bf9e4e30f353 ("x86/mm: use pagetable_free()"), switched from freeing non-boot page tables through __free_pages() to pagetable_free(). However, the function is also called to free vmemmap pages. Given that vmemmap pages are not page tables, already the page_ptdesc(page) is wrong. But worse, pagetable_free() calls: __free_pages(page, compound_order(page)); Since vmemmap pages are not compound pages (see vmemmap_alloc_block()) -- except for HVO, which doesn't apply here -- only first page of a PMD-sized vmemmap page is freed, leaking the other ones. Fix it by properly decoupling pagetable and vmemmap freeing. free_pagetable() no longer has to mess with SECTION_INFO, as only the vmemmap is marked like that in register_page_bootmem_memmap(). The indentation in remove_pmd_table() is messed up. Fix that while touching it. Bootmem info handling will soon be fixed up. For now, handle it similar to free_pagetable(), just avoiding the ifdef. [ dhansen: changelog munging. More imperative voice ]

In the Linux kernel, the following vulnerability has been resolved: regulator: scmi: fix of_node refcount leak in scmi_regulator_probe() scmi_regulator_probe() calls of_find_node_by_name() which takes a reference on the returned device node. On the error path where process_scmi_regulator_of_node() fails, the function returns without calling of_node_put() on the child node, leaking the reference. Add of_node_put(np) on the error path to properly release the reference.

In the Linux kernel, the following vulnerability has been resolved: perf/aux: Fix page UAF in map_range() map_range() reads rb->aux_pages[], rb->aux_nr_pages and rb->aux_pgoff via perf_mmap_to_page() while holding only event->mmap_mutex. Those fields are serialized by rb->aux_mutex, and mmap_mutex is per event. Thus, two events sharing one rb via PERF_EVENT_IOC_SET_OUTPUT can race rb_alloc_aux() with map_range(), leading to a page-UAF scenario as follows: CPU 0 CPU 1 ===== ===== rb_alloc_aux() map_range() [1]: allocate rb->aux_pages[0] [2]: rb->aux_nr_pages++ [3]: perf_mmap_to_page() returns rb->aux_pages[0] [4]: map it as VM_PFNMAP [5]: rb->aux_pgoff = 1 munmap the page [6]: free rb->aux_pages[0] Pages mapped as VM_PFNMAP have no refcount protection, so CPU 1 holds a mapping to a freed physical frame. Fix this by taking rb->aux_mutex across the page walk in map_range().

In the Linux kernel, the following vulnerability has been resolved: tracing: Prevent out-of-bounds read in glob matching String event fields are not necessarily NUL-terminated, so the filter predicate functions (filter_pred_string(), filter_pred_strloc() and filter_pred_strrelloc()) pass the field length to the regex match callbacks, and the length-aware matchers honour it. regex_match_glob() was the exception: it ignored the length and called glob_match(), which scans the string until it hits a NUL byte. Some string fields are not NUL-terminated. One example is the dynamic char array of the xfs_* namespace tracepoints, which is copied without a trailing NUL. For such a field, glob matching reads past the end of the event field, causing a KASAN slab-out-of-bounds read in glob_match(), reached via regex_match_glob() and filter_match_preds() from the xfs_lookup tracepoint. Add a length-bounded glob_match_len() and use it from regex_match_glob() so glob matching always stops at the field boundary. The matching loop is factored into a shared helper so glob_match() keeps its behaviour.

In the Linux kernel, the following vulnerability has been resolved: NFSv4: include MAY_WRITE in open permission mask for O_TRUNC POSIX requires write permission to truncate a file, so an open() that specifies O_TRUNC must be authorized for write access regardless of the O_ACCMODE access mode. nfs_open_permission_mask() builds the access mask passed to nfs_may_open(), which is the local authorization gate for OPENs the client serves itself from a cached write delegation via the can_open_delegated() path in nfs4_try_open_cached(). The mask is derived from O_ACCMODE alone, so an open(O_RDONLY | O_TRUNC) against a file the caller cannot write requests only MAY_READ and passes the local check. The OPEN is then satisfied locally and the truncation is issued to the server as a SETATTR(size=0) over the delegation stateid, which the server accepts under standard write-delegation semantics. POSIX requires that this open fail with EACCES. Include MAY_WRITE in the mask whenever O_TRUNC is set so the local check matches the access the server would have enforced.

In the Linux kernel, the following vulnerability has been resolved: module: decompress: check return value of module_extend_max_pages() module_extend_max_pages() calls kvrealloc() internally and returns -ENOMEM on allocation failure. The return value is never checked. If the initial allocation fails, info->pages remains NULL and info->max_pages remains 0. Subsequent calls to module_get_next_page() will attempt to dynamically grow the array by calling module_extend_max_pages(info, 0) since info->used_pages is 0. This results in kvrealloc(NULL, 0) returning ZERO_SIZE_PTR, which is treated as a success, leading to a dereference of ZERO_SIZE_PTR and a kernel oops. Fix: add the missing error check after module_extend_max_pages() and return immediately on failure. This matches the pattern used by every other kvrealloc() caller in the module loading path. [Sami: Corrected the analysis in the commit message.]

In the Linux kernel, the following vulnerability has been resolved: exfat: bound uniname advance in exfat_find_dir_entry() In exfat_find_dir_entry(), each TYPE_EXTEND (file name) entry advances the output pointer by a fixed amount while the loop guard only tracks the accumulated name length: if (++order == 2) uniname = p_uniname->name; else uniname += EXFAT_FILE_NAME_LEN; len = exfat_extract_uni_name(ep, entry_uniname); name_len += len; unichar = *(uniname+len); *(uniname+len) = 0x0; uniname grows by EXFAT_FILE_NAME_LEN (15) per name entry, but name_len grows only by the actual extracted length, which is shorter when a name fragment contains an early NUL. The only guard is `name_len >= MAX_NAME_LENGTH`, so a crafted directory with many short name fragments lets uniname run far past the p_uniname->name[MAX_NAME_LENGTH + 3] buffer while name_len stays small, causing an out-of-bounds read and write at *(uniname+len). The sibling extractor exfat_get_uniname_from_ext_entry() already stops on a short fragment (the lockstep `len != EXFAT_FILE_NAME_LEN` guard added in commit d42334578eba ("exfat: check if filename entries exceeds max filename length")); exfat_find_dir_entry() never got the equivalent. Track the per-entry write offset as a count and reject a fragment once the offset, or the offset plus the extracted length, would exceed MAX_NAME_LENGTH, before forming the output pointer.

In the Linux kernel, the following vulnerability has been resolved: mm: page_ext: add count limit to page_ext_iter_next to prevent invalid PFN access The page_ext iteration API does not validate if the PFN still belongs to a valid section while advancing the iterator. When dynamically adding memory in the hotplug path, it can lead to a NULL pointer dereference during page_ext_lookup at the boundary of the last valid section when iterator count equals __pgcount. The for_each_page_ext() macro calls page_ext_iter_next() as its loop increment. for_each_page_ext() does a "__page_ext = page_ext_iter_next(&__iter)" at the end. This causes page_ext_iter_next() to increment iter->index past __pgcount and call page_ext_lookup(start_pfn + __pgcount). During memory hotplug (online), the PFN at start_pfn + __pgcount may belong to a section that has not yet been initialized, causing page_ext_lookup() to trigger a NULL pointer dereference. [ 14.555124][ T846] Call trace: [ 14.555125][ T846] lookup_page_ext+0x6c/0x108 (P) [ 14.555127][ T846] page_ext_lookup+0x30/0x3c [ 14.555129][ T846] __reset_page_owner+0x11c/0x260 [ 14.571201][ T846] __free_pages_ok+0x5e8/0x8e0 [ 14.571204][ T846] __free_pages_core+0x78/0xf0 [ 14.571206][ T846] generic_online_page+0x14/0x24 [ 14.597782][ T846] online_pages+0x178/0x30c [ 14.597784][ T846] memory_block_change_state+0x284/0x32c [ 14.597787][ T846] memory_subsys_online+0x4c/0x64 [ 14.597789][ T846] device_online+0x88/0xb0 [ 14.597791][ T846] online_memory_block+0x30/0x40 [ 14.597793][ T846] walk_memory_blocks+0xac/0xe8 [ 14.597794][ T846] add_memory_resource+0x280/0x298 [ 14.656161][ T846] add_memory+0x60/0x98 Move the iteration boundary enforcement inside the iterator functions, so callers cannot inadvertently access beyond the requested range.

In the Linux kernel, the following vulnerability has been resolved: mm: do file ownership checks with the proper mount idmap Ever since idmapped mounts were introduced, inode ownership checks (for side-channel protection) in mincore() and madvise(MADV_PAGEOUT) were done against the nop_mnt_idmap, which completely ignores the file's mount's idmap. This results in odd edgecases like: 1) mount/bind-mount with an idmap userA:userB:1 2) userB runs an owner_or_capable() check on file that is owned by userA on-disk/in-memory, but owned by userB after idmap translation 3) owner_or_capable() mysteriously fails as the correct idmap wasn't supplied In the case of mincore/madvise MADV_PAGEOUT, this is usually benign, because file_permission(file, MAY_WRITE) will probably succeed, as it uses the proper idmap internally, but it does not need to be the case on e.g a 0444 file where even the owner itself doesn't have permissions to write to it. Since this is clearly not trivial to get right, introduce a file_owner_or_capable() that can carry the correct semantics, and switch the various users in mm to it. The issue was found by manual code inspection & an off-list discussion with Jan Kara.

In the Linux kernel, the following vulnerability has been resolved: iommufd: Use sizeof(*hdr) instead of sizeof(hdr) in veventq read The bound-check in iommufd_veventq_fops_read() for the normal vEVENT path uses sizeof(hdr) where the surrounding code uses sizeof(*hdr): if (!vevent_for_lost_events_header(cur) && sizeof(hdr) + cur->data_len > count - done) { hdr is declared as struct iommufd_vevent_header *, so sizeof(hdr) evaluates to the size of the pointer. Surrounding code uses sizeof(*hdr) consistently: if (done >= count || sizeof(*hdr) > count - done) { ... if (copy_to_user(buf + done, hdr, sizeof(*hdr))) { ... done += sizeof(*hdr); struct iommufd_vevent_header is currently 8 bytes (two __u32 fields, flags and sequence), so on 64-bit (sizeof(void *) == 8) the two expressions happen to be equal and the check works as intended. On 32-bit (sizeof(void *) == 4) the check under-counts the header by 4 bytes: a vEVENT whose data_len causes 8 + cur->data_len to exceed count - done while 4 + cur->data_len does not will pass the check, then the loop will copy_to_user 8 bytes of header followed by data_len bytes of payload, writing past the user-supplied buffer. It is also a latent bug for any future expansion of struct iommufd_vevent_header beyond sizeof(void *) on 64-bit; the check should not depend on the type happening to match the host pointer width. Use sizeof(*hdr) to match the rest of the function and the actual amount that will be copied.

In the Linux kernel, the following vulnerability has been resolved: iommufd: Move vevent memory allocation outside spinlock The veventq memory allocation happens inside the spinlock. Given its depth is decided by the user space, this leaves a vulnerability, where userspace can allocate large queues to exhaust atomic memory reserves. Move the allocation outside the spinlock and use GFP_NOWAIT, which can fail fast under memory pressure without dipping into the GFP_ATOMIC reserves or direct-reclaiming from the threaded IRQ handler. On allocation failure, queue the lost_events_header (so userspace learns of the drop) and return -ENOMEM so the caller learns of the kernel-side memory pressure. This is intentionally distinct from the queue-overflow path, which also queues the lost_events_header but returns 0: a full queue is an expected userspace-pacing condition rather than a kernel error. A subsequent change will cap the upper bound of the veventq_depth.

In the Linux kernel, the following vulnerability has been resolved: iommufd: Set veventq_depth upper bound iommufd_veventq_alloc() accepts any !0 veventq_depth from userspace, with an upper bound at U32_MAX. This leaves a vulnerability where userspace can allocate excessively large queues to exhaust kernel memory reserves. Cap the veventq_depth (maximum number of entries) to 1 << 19, matching the maximum number of entries in the SMMUv3 EVTQ (the largest use case today).

In the Linux kernel, the following vulnerability has been resolved: iommufd: Break the loop on failure in iommufd_fault_fops_read() On a copy_to_user() failure inside the inner list_for_each_entry, only the inner loop breaks; the outer while re-fetches the just-restored fault group and retries the failing copy_to_user() forever, spinning the reader at 100% CPU with fault->mutex held. Check rc after the inner loop and break the outer while as well.

In the Linux kernel, the following vulnerability has been resolved: iommufd: Set upper bounds on cache invalidation entry_num and entry_len iommufd_hwpt_invalidate() takes a user-controlled entry_num and entry_len, each bounded only by U32_MAX. An entry_len beyond the kernel's struct size makes the copy helper verify the extra bytes are zero, scanning that excess in one uninterruptible pass; a multi-gigabyte value over zeroed user memory trips the soft-lockup watchdog. A large entry_num is the other half, driving the backend invalidation loop with no reschedule. The VT-d nested handler, for one, copies each entry and flushes caches per iteration, pinning the CPU on a non-preemptible kernel. Cap both in the ioctl. entry_len is held under PAGE_SIZE, above any request struct, and entry_num under 1 << 19, the order of a hardware invalidation queue and well beyond any real batch, bounding the per-call loop length.

In the Linux kernel, the following vulnerability has been resolved: KVM: arm64: nv: Avoid dereferencing NULL VNCR pseudo-TLB VNCR TLB invalidation occurs from MMU notifiers or TLBI instructions, and either can race against a vcpu not being onlined yet (no pseudo-TLB allocated). Similarly, the TLB might be invalid, and the invalidation should be skipped in this case. Both kvm_invalidate_vncr_ipa() and kvm_invalidate_vncr_va() are expected to perform the same checks, except that the latter doesn't check for the allocation and blindly dereferences the pointer. Solve this by introducing a new iterator built on top of the usual kvm_for_each_vcpu() that checks for both of the above conditions, and convert the two users to it.