Only Meltdown is specifically an Intel vulnerability / design flaw.
update: it seems AMD is mostly resilient to Spectre. It's not clear why that would be the case. But according to AMD:
(from early January, now replaced, see update 2 below)
Differences in AMD architecture mean there is a near zero risk of exploitation of this variant. Vulnerability to Variant 2 (Branch Target Injection) has not been demonstrated on AMD processors to date.
Note that it's only a "near zero", unlike the "zero" for Meltdown. (And that the "bounds check bypass" Spectre v1 attack (conditional direct branch before an array access) still needs to be fixed in software.)
update 2: AMD may have underestimated Spectre, because they are still releasing microcode updates, and their web site now says:
GPZ Variant 2 (Branch Target Injection or Spectre) is applicable to AMD processors.
- While we believe that AMD’s processor architectures make it difficult to exploit Variant 2, we continue to work closely with the industry on this threat.
It's still not clear what microarchitectural design feature / difference makes AMD CPUs any more resilient than Intel; they do still have a large out-of-order window and AFAIK are not significantly more limited in what can be in flight at once. But their branch predictors are designed differently, and are maybe harder to mis-train or are somehow insulated across privilege boundaries? But the latter would make them immune to Spectre v1 as well, and AMD CPUs are just as vulnerable to the bounds-check bypass version of Spectre.
Spectre (v2) affects any "normal" CPU design with branch prediction + speculative execution for indirect branches. In Spectre, the attacker primes / trains the branch predictor to mispredict the target of an indirect branch in kernel code or in another process. Until the mis-predict is discovered, the privileged process will speculatively execute some instructions that change the micro-architectural state of the machine in a way that depends on secret data.
(Then the attacker uses a cache-timing attack to read that microarchitectural state, the same as in Meltdown. When attacking the kernel, you'd ideally choose a "gadget" that uses the secret data to index an array your attacking user-space code also has read privilege for. If you can't, you might have to use a cache-eviction timing attack to look for which set of cache had a line evicted instead of looking for a line that became hot.)
The key to Meltdown is that Intel CPUs (but apparently not ARM or AMD) don't squash under-privileged TLB hits. The load executes, and so do following instructions, and it only actually faults when the faulting load tries to retire. This allows unprivileged code to directly cause changes in the microarchitectural state itself, based on data it doesn't have permission to read, instead of (Spectre) tricking a privileged process into doing it.
;; run this in user-space
;; (and suppress or catch the fault somehow so you can do it quickly/repeatedly)
;; clflush all the cache lines in local_array (which you have permission to read)
;; create a long delay before following instructions can retire, but with few uops so OoO exec can see past it
times 30 sqrtpd xmm0, xmm0 ; high latency per uop
movzx eax, byte [kernel_byte] ; eventually faults, but OoO exec continues first
shl eax, 12
mov eax, [local_array + rax] ; the cache-line this touches depends on the secret data
;; after the movzx faults, the cache-line touched by the mov will still be hot
; then check which cache line of local_array is already hot.
(And BTW, I emailed the authors of the Meltdown paper about this. They were in a rush to get it published by the public disclosure date which apparently moved up by 1 week unexpectedly. They are planning to clarify their section on the microarchitectural details to make it clear that this is the real design flaw that enables Meltdown.)
(Meltdown depends on the kernel keeping its pages mapped into the virtual address space of the process, but with a bit set in the page table entries that flag them as kernel-only mappings. This makes system calls and interrupts cheaper, because they don't have to modify the page tables to allow access to kernel memory: the CPU just starts allowing those kernel-only mappings to work when running in ring 0. See this Q&A for a diagram of the x86-64 page-table entry format. The U/S bit (user/supervisor) is the one that controls whether a mapping is kernel-only or not.)
The fault that's eventually raised is either handled (catch SIGSEGV), or suppressed (run the transient sequence inside a TSX (transactional memory) transaction, or as a result of mis-speculation). So in Meltdown, branch mispredicts are only relevant as part of making the attack efficient and reliable by suppressing the fault on the load from a kernel virtual address. Speculative execution past a faulting load is key, though, even for dependent instruction. (Not just for independent instructions executed before a load-address is ready or anything that simple).
Presumably AMD's load execution units / TLB are designed differently, with privilege checking for loads applied earlier or differently. Either a load from a virtual address with a kernel-only mapping is treated the same as a load from an unmapped page, or the physical page bits used for speculative execution are set to all-zero or all-one or something. Or maybe it just squashes the load without triggering a page-walk (like a load from an unmapped page would).
Note that x86 doesn't require TLB invalidation when changing a page table entry from invalid to valid, so the address-translation hardware isn't allowed to do "negative" caching. Or if it does, it would have to be coherent with page table writes that make previously-invalid entries valid. Intel CPUs do a kind of TLB shootdown to maintain coherence for TLB entries that were only speculatively loaded, going beyond what the x86 manuals require to avoid breaking old code that existed before the current TLB-invalidation rules were published (e.g. Win95 through ME).
The point is, a different microarchitectural design choice could totally block the thing that Meltdown depends on. And such design choices are plausible on their own, not specifically to avoid Meltdown.
A related question: why does Intel's delayed permission check design choice make sense?
Until Meltdown / Spectre, CPU designers were only worried about making sure memory protection applied to the architectural state (non-speculative values in architectural registers, not physical registers used by out-of-order execution). i.e. this side-channel wasn't on anyone's radar. The results of instruction execution don't become architectural state until retirement, so that's the point when everything has to be correct (in pre-Meltdown thinking).
As a CPU designer, you want everything to execute as efficiently as possible, with as few special cases as possible. Or especially, with special cases in as few components as possible. It's a lot simpler for the overall design if an execution unit can never stall, so that a pipelined section of logic doesn't need any flow-control, just simply always accept one new input per clock.
It looks like Intel's load execution units don't have a way to squash the load on a TLB hit. The TLB itself is a CAM (Content-Addressable Memory). Managing TLB entries with speculative (prefetch) page-walks involves some more active hardware, but the TLB itself has to be fast to support 3 lookups per clock (from ports 2, 3, and 7).
Most code doesn't page-fault by trying to load from kernel addresses, so optimizing for that case wasn't a consideration. If such loads are seen by out-of-order execution, they usually happen as a result of mis-speculation (running a load instruction with the wrong data in register). (Not necessarily maliciously induced mis-speculation (Spectre), just the regular kind from imperfect branch prediction.) Not doing anything about a TLB-hit faulting load until retirement is a good design decision if most of the time it never reaches retirement because a branch mispredict or other mis-speculation is detected earlier in in-order retirement. Wasting load bandwidth and causing cache pollution is questionable, but (other than Meltdown attacks) this is probably a pretty rare case so keeping the hardware simple was the most valuable thing.
Usually page-faults happen because of access to an address which isn't mapped at all: the page table entry is marked Invalid, not just kernel-only. Or even a higher level of the 4-level nested page tables is invalid, e.g. the page-directory entry. As mentioned earlier, negative caching isn't architecturally allowed (and AFAIK isn't done with snooping for correctness either), so such PTEs will never appear in the TLB. A page-fault for an unmapped page will be raised only after a page-walk (which finds the mapping for that page doesn't exist). x86 has dedicated page-walk hardware so the page-table loads can happen in the background while the execution units run other uops. (Skylake even has two HW page-walk units). But anyway,
So a page-fault from trying to read a kernel-only mapping is a special case, very different microarchitecturally from the unmapped-page case. (It's actually similar to trying to store into a read-only mapping, which presumably also has delayed faulting. Stores don't become globally visible right away; the store buffer makes speculative store execution possible by keeping them private until after retirement, at which point they can commit to L1D).
How could Intel fix Meltdown in future hardware?
Fixing Meltdown is relatively easy (compared to Spectre), although it probably can't be done with a microcode update. As well as setting a fault-if/when-this-reaches-retirement bit on the uop, a TLB lookup could gate the page-address bits (to all ones) with the privilege-check. e.g. a load in user-space from any kernel page could micro-architecturally execute as a load from the very top physical page. (And systems with less than the max amount of RAM wouldn't have any physical RAM at that physical address.)
Or a failed privilege check could maybe still allow the load to happen microarchitecturally, but mask the result to all-zero in the load port. (Remember, the Meltdown problem isn't that an unprivileged load can bring kernel data into cache, it's that the secret data load result can be used to make another load with a data-dependent address. Continuing speculative execution with a zero result for any under-privileged load that hits in the TLB wouldn't allow any data-dependent microarchitectural effects).
The TLB lookup happens in parallel with indexing the VIPT L1D cache, but the TLB result is needed for the tag-check part of an L1D cache load. So requiring the TLB lookup result is already needed to select the right way from the set indexed by address bits 6 to 11. (L1D is 8-way set associative). So also requiring the TLB permission check to be ready this early shouldn't introduce any extra latency. Masking with the permission result would introduce one extra gate-delay, but one clock cycle has time for many gate delays. (e.g. 64-bit
add latency is only 1 cycle, 64-bit
imul latency is only 3 cycles on Intel Sandybridge-family. http://agner.org/optimize/).
You could even design it so a load that would fault (if it reaches retirement) doesn't complete execution at all, and no load result is forwarded to dependent instructions. Maybe even squash it so it doesn't send a memory request down the hierarchy towards outer caches if it wasn't in L1D cache. (i.e. don't even check L1D for it). This might be a more tricky design.
This design would make it impossible(?) for a user-space process to even detect whether a kernel address was mapped or not by any kind of timing attack, because the micro-architectural effect of a load to a kernel-only mapping would be the same as to an unmapped page. (Hmm, actually only the same if it also triggered a page-walk.)
This might be valuable to stop processes defeating KASLR, if the kernel uses anything smaller than 1G hugepages to map any of its own memory.
Further reading about CPU internals: