Hardware Translation Coherence for Virtualized Systems

01/25/2017 ∙ by Zi Yan, et al. ∙ Rutgers University 0

To improve system performance, modern operating systems (OSes) often undertake activities that require modification of virtual-to-physical page translation mappings. For example, the OS may migrate data between physical frames to defragment memory and enable superpages. The OS may migrate pages of data between heterogeneous memory devices. We refer to all such activities as page remappings. Unfortunately, page remappings are expensive. We show that translation coherence is a major culprit and that systems employing virtualization are especially badly affected by their overheads. In response, we propose hardware translation invalidation and coherence or HATRIC, a readily implementable hardware mechanism to piggyback translation coherence atop existing cache coherence protocols. We perform detailed studies using KVM-based virtualization, showing that HATRIC achieves up to 30 energy benefits, for per-CPU area overheads of 2 benefits on systems running Xen and find up to 33

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1 Introduction

As the computing industry designs systems for big-memory workloads, systems architects have begun embracing heterogeneous memory architectures. For example, Intel is integrating high-bandwidth on-package memory in its Knight’s Landing chip, and 3D Xpoint memory in several products [1]. AMD and Hynix are releasing High-Bandwidth Memory or HBM [2, 3]. Similarly, Micron’s Hybrid Memory Cube [4, 5] and byte-addressable persistent memories [6, 7, 8, 9] are quickly gaining traction. Vendors are combining these high-performance memories with traditional high-capacity and low-cost DRAM, prompting research on heterogeneous memory architectures [9, 10, 2, 11, 12, 13, 14, 15].

Fundamentally, heterogeneous memories are dependent on the concept of page remapping to migrate data between diverse memory devices for good performance. Page remapping is not a new concept – OSes have long used it to migrate physical pages to defragment memory and create superpages [16, 17, 18, 19], to migrate pages among NUMA sockets [20, 21], and to deduplicate memory by enabling copy-on-write optimizations [22, 23, 24]. However, while page remappings were used sparingly in those scenarios, they are likely to be used more frequently for heterogeneous memories. This is because page remapping is essential to adapt data placement to the memory access patterns of workloads, and to harness the performance and energy potential of memories with different latency, bandwidth, and capacity characteristics. Consequently, developers at IBM and Redhat are already deploying Linux patchsets to enable page remapping amongst coherent heterogeneous memory devices [25, 26, 27].

Unfortunately, these efforts face an obstacle – the high performance and energy penalty of page remapping. There are two components to this cost. The first is the overhead of copying data. The second is the cost of translation coherence. When privileged software remaps a physical page, it has to update the corresponding virtual-to-physical page translation in the page table. Translation coherence is the means by which caches dedicated to translations (e.g., Translation Lookaside Buffers or TLBs [28, 29, 30, 31], etc.) are kept up to date with the latest page table mappings.

Past work has shown that translation coherence overheads can easily consume 10-30% of system performance [10, 32, 33]. These overheads are even more alarming on virtualized systems, which are used in the server and cloud settings expected to be early adopters of heterogeneous memories. We are the first to show that as much as 40% of their runtime can be wasted on translation coherence. The key culprit is virtualization’s use of multiple page tables. Architectures with hardware assists for virtualization like Intel VT-x and AMD-V use a guest page table to map guest virtual pages to guest physical pages, and a nested page table to map guest physical pages to system physical pages. Changes to the guest page table and in particular, the nested page table, prompt expensive translation coherence activity.

The problem of coherence is not restricted to translation mappings. In fact, the systems community has studied problems posed by cache coherence for several decades [34] and has developed efficient hardware cache coherence protocols [35]. What makes translation coherence challenging is that unlike cache coherence, it relies on cumbersome software support. While this may have sufficed in the past when page remappings were used relatively infrequently, they are problematic for heterogeneous memories where page remapping is more frequent. Consequently, we believe that there is a need to architect better support for translation coherence. In order to understand what this support should constitute, we list three attributes desirable for translation coherence.

1⃝ Precise invalidation: Processors use several hardware translation structures – TLBs, MMU caches [36, 37], and nested TLBs (nTLBs) [38] – to cache portions of the page table(s). Ideally, translation coherence should invalidate the translation structure entries corresponding to remapped pages, rather than flushing all the contents of these structures.

2⃝ Precise target identification: The CPU running privileged code that remaps a page is known as the initiator. An ideal translation coherence protocol would allow the initiator to identify and alert only CPUs whose TLBs, MMU caches, and nTLBs actually cache the remapped page’s translation. By restricting coherence messages to only these targets, other CPUs remain unperturbed by coherence activity.

3⃝ Lightweight target-side handling: Target CPUs should invalidate their translation structures and relay acknowledgment responses to the initiator quickly, without excessively interfering with workloads executing on the target CPUs.

Unfortunately, translation coherence meets none of these goals today. Consider, for example, changes to the nested page table. Further, consider 1⃝; when hypervisors change a nested page table entry, they track guest physical and system physical page numbers, but not the guest virtual page. Unfortunately, as we describe in Sec. 2, translation structures on architectures like x86-64 permit invalidation of individual entries only if their guest virtual page is known. Consequently, hypervisors completely flush all translation structures, even when only a single page is remapped. This degrades performance since virtualized systems need expensive two-dimensional page table walks to re-populate the flushed structures [22, 38, 39, 40, 41, 42, 43, 28, 29].

Current translation coherence protocols also fail to achieve 2⃝. Hypervisors track the subset of CPUs that a guest VM runs on but cannot (easily) identify the CPUs used by a process within the VM. Therefore, when the hypervisor remaps a page, it conservatively initiates coherence activities on all CPUs that may potentially have executed any process in the guest VM. While this does spare CPUs that never execute the VM, it needlessly flushes translation structures on CPUs that execute the VM but not the process.

Finally, 3⃝ is also not met. Initiators currently use expensive inter-processor interrupts (on x86) or tlbi instructions (on ARM, Power) to prompt VM exits on all target CPUs. Translation structures are flushed on a VM re-entry. VM exits are particularly detrimental to performance, interrupting the execution of target-side applications [38, 44].

We believe that the solution to these problems is to implement translation coherence in hardware. This view is inspired by prior work on UNITD [32], which showcased the potential of hardware translation coherence. Unfortunately, UNITD is energy inefficient and, like other recent proposals [10, 33], cannot support virtualized systems. In response, we propose hardware translation invalidation and coherence or HATRIC, a hardware mechanism to tackle these problems and meet 1⃝-3⃝. HATRIC extends translation structure entries with coherence tags (or co-tags) storing the system physical address where the translation entry resides (not to be confused with the physical address stored in the page table). This solves 1⃝, since translation structures can now be identified by the hypervisor without knowledge of the guest virtual address. HATRIC exposes co-tags to the underlying cache coherence protocol, achieving 2⃝ and 3⃝.

We evaluate HATRIC for a forward-looking virtualized system with a high-bandwidth die-stacked memory and a slower off-chip memory. HATRIC drastically reduces translation coherence overheads, improving performance by 30%, saving as much as 10% of energy, while adding less than 2% of CPU area. Overall, our contributions are:

  • We perform a characterization study to quantify the overheads of translation coherence on hypervisor-managed die-stacked memory. While we focus on KVM in this paper, we have also studied Xen and quantified its overheads.

  • We design HATRIC to subsume translation coherence in hardware by piggybacking on, without fundamentally changing, existing cache coherence protocols. HATRIC goes beyond UNITD [32] by a⃝ accommodating translation coherence for both bare-metal and virtualized scenarios; b⃝ extending coherence to not just TLBs, but also MMU caches and nTLBs; c⃝ and achieving better energy efficiency.

  • We perform several studies that illustrate the benefits of HATRIC’s design decisions. Further, we discuss HATRIC’s advantages over purely software approaches to mitigate translation coherence issues.

Overall, HATRIC is efficient and versatile. While we mostly focus on the particularly arduous challenges of translation coherence due to nested page table changes, HATRIC is applicable to guest page tables and non-virtualized systems.

2 Background

We begin by presenting an overview of the key hardware and software structures involved in page remapping. Our discussion focuses on x86-64 systems. Other architectures are broadly similar but differ in some low-level details.

2.1 HW and SW Support for Virtualization

Virtualized systems accomplish virtual-to-physical address translation in one of two ways. Traditionally, hypervisors have used shadow page tables to map guest virtual pages (GVPs) to system physical pages (SPPs), keeping them synchronized with guest OS page tables [40]. However, the overheads of page table synchronization can often be high [41]. As a result, most modern systems now use two dimensional page tables instead. Figure 1 illustrates two-dimensional page table walks (see past work for more details [22, 36, 38, 40, 45, 46, 42]). Guest page tables map GVPs to guest physical pages (GPPs). Nested page tables map GPPs to SPPs. x86-64 systems use 4-level forward mapped radix trees for both page tables [22, 38, 46, 45]. We refer to these as levels 4 (the root level) to 1 (the leaf level) as per recent work [38, 37, 36]. When a process running in a guest VM makes a memory reference, its GVP must be translated to an SPP. Consequently, the guest CR3 register is combined with the requested GVP (not shown in the picture) to deduce the GPP of level 4 of the guest page table (shown as GPP Req.). However, to look up the guest page table (gL4-gL1), the GPP must be converted into the SPP where the page table actually resides. Therefore, we first use the GPP to look up the nested page tables (nL4-nL1), to find SPP gL4. Looking up gL4 then yields the GPP of the next guest page table level (gL3). The rest of the page table walk proceeds similarly, requiring 24 memory references in total. This presents a performance problem as the number of references is significantly more than the 4 references needed for non-virtualized systems. Further, the references are entirely sequential. CPUs use three types of translation structures to accelerate this walk:

Figure 1: Two-dimensional page table walks for virtualized systems. Nested page tables are represented by boxes and guest page tables are represented by circles. Each page table’s levels from 4 to 1 are shown. We show items cached by MMU caches and nTLBs. TLBs (not shown) cache translations from the requested guest virtual page (GVP) to the requested system physical page (SPP).

a⃝ Private per-CPU TLBs cache the requested GVP to SPP mappings, short-circuiting the entire walk. TLB misses trigger hardware page table walkers to look up the page table.

b⃝ Private per-CPU MMU caches store intermediate page table information to accelerate parts of the page table walk [36, 37, 38]. There are two flavors of MMU cache. The first is a page walk cache and is implemented in AMD chips [37, 38]. Figure 1 shows the information cached in page walk caches. Page walk caches are looked up with GPPs and provide SPPs where page tables are stored. The second is called a paging structure cache and is implemented by Intel [37, 36]. Paging structure caches are looked up with GVPs and provide the SPPs of page table locations. Paging structure caches generally perform better, so we focus mostly on them [37, 36].

c⃝ Private per-CPU nTLBs short-circuit nested page table lookups by caching GPP to SPP translations [38]. Figure 1 shows the information cached by nTLBs.

Concomitantly, CPUs cache page table information in private L1 (L2, etc.) caches and the shared last-level cache (LLC). The presence of separate private translation caches poses coherence problems. While standard cache coherence protocols ensure that page table entries in private L1 caches are coherent, there are no such guarantees for TLBs, MMU caches, and nTLBs. Instead, privileged software keeps translation structures coherent with data caches and one another.

2.2 Page Remapping in Virtualized Systems

We now detail the ways in which a virtualized system can trigger coherence activity in translation structures. All page remappings can be classified by the data they move, and the software agent initiating the move.

Remapped data: Systems may remap a page storing (i) the guest page table; (ii) the nested page table; or (iii) non-page table data. Most remappings are from (iii) as they constitute most memory pages. We have found that less than 1% of page remappings correspond to (i)-(ii). We therefore highlight HATRIC’s operation using (iii); nevertheless, HATRIC also implicitly supports the first two cases.

Remapping initiator: Pages can be remapped by (i) a guest OS; or (ii) the hypervisor. When a guest OS remaps a page, the guest page table changes. Past work achieves low-overhead guest page table coherence with relatively low-complexity software extensions [47]. Unfortunately, there are no such workarounds to mitigate the translation coherence overheads of hypervisor-initiated nested page table remappings. For these reasons, cross-VM memory deduplication [22, 48] and page migration between NUMA memories on multi-socket systems [49, 50, 51] are known to be expensive. In the past, such overheads may have been mitigated by using these optimizations sparingly. However, nested page table remappings become frequent with heterogeneous memories, making hypervisor-initiated translation coherence problematic.

3 Shortcomings of Current Translation Coherence Mechanisms

Our goal is to ensure that translation coherence does not impede the adoption of heterogeneous memories. We study forward-looking die-stacked DRAM as an example of an important heterogeneous memory system. Die-stacked memory uses DRAM stacks that are tightly integrated with the processor die using high-bandwidth links like through-silicon vias, or silicon interposers [10, 52]. Die-stacked memory is expected to be useful for multi-tenant and rack-scale computing where memory bandwidth is often a performance bottleneck, and will require a combination of application, guest OS, and hypervisor management [10, 53, 54, 55]. We take the first steps towards this, by showing the problems posed by translation coherence on hypervisor management.

3.1 Translation Coherence Overheads

We quantify translation coherence overheads on a die-stacked system that is virtualized with KVM. We modify KVM to page between the die-stacked and off-chip DRAM. Since ours is the first work to consider hypervisor management of die-stacked memory, we implement a variety of paging policies. Rather than focusing on developing a single “best” policy, our objective is to show that current translation coherence overheads are so high that they curtail the effectiveness of practically any paging policy.

Our paging mechanisms extend prior work that explores basic software-guided die-stacked DRAM paging [10]. When off-chip DRAM data is accessed, there is a page fault. KVM then migrates the desired page into an available die-stacked DRAM physical page frame. The GVP and GPP remain unchanged, but KVM changes the SPP and hence, its nested page table entry. This triggers translation coherence.

We run our modified KVM on the detailed cycle-accurate simulator described in Sec. 5. Like prior work [10], we model a system with 2GB of die-stacked DRAM with 4 the memory bandwidth of a slower off-chip 8GB DRAM. This is a total of 10GB of addressable DRAM. Further, we model 16 CPUs based on Intel’s Haswell architecture.

Figure 2: Performance of no-hbm (no die-stacked DRAM), inf-hbm (data always in die-stacked DRAM), curr-best (best die-stacked DRAM paging policy with current software translation coherence overheads), and achievable (best achievable paging policy, assuming no translation coherence overheads). All data is normalized to no-hbm runtime.

Figure 2 quantifies the performance of hypervisor-managed die-stacked DRAM, and translation coherence’s impact on it. We normalize all performance numbers to the runtime of a system with only off-chip DRAM and no high-bandwidth die-stacked DRAM (no-hbm). Further, we show an unachievable best-case scenario where all data fits in an infinite-sized die-stacked memory (inf-hbm). After profiling several paging strategies (evaluated in detail in Sec. 6), we plot the best-performing ones with the curr-best bars. These results assume cumbersome software translation coherence mechanisms. In contrast, the achievable bars represent the potential performance of the best paging policies with zero-overhead (and hence ideal) translation coherence.

Figure 2 shows that unachievable infinite die-stacked DRAM can improve performance by 25-75% (inf-hbm versus no-hbm). Unfortunately, the current “best” paging policies we achieve in KVM (curr-best) fall far short of the ideal inf-hbm case. Translation coherence overheads are a big culprit – when these overheads are eliminated in achievable, system performance comes within 3-10% of the case with infinite die-stacked DRAM capacity (inf-hbm). In fact, Figure 2 shows that translation coherence overheads can be so high that they can prompt die-stacked DRAM to counterintuitvely worsen performance. For example, data caching and tunkrank actually suffer 23% and 10% performance degradations in curr-best, respectively, despite using high-bandwidth die-stacked memory. Though omitted to save space, we have also profiled the Xen hypervisors and found similar trends (presented in Sec. 6). Overall, translation coherence overheads threaten the use of die-stacked, and indeed any heterogeneous, memory.

3.2 Page Remapping Anatomy

Figure 3: Sequence of operations associated with a page unmap. Initiator to target IPIs are shown in blue 1⃝, VM exits are shown in green 2⃝, and translation structure flushes are shown in black 3⃝.

We now shed light on the sources of overheads from translation coherence. While we use page migration between off-chip and die-stacked DRAM as our driving example, the same mechanisms are used today to migrate pages between NUMA memories, or to defragment memory, etc.

When a VM is configured, KVM assigns it virtual CPU threads or vCPUs. Figure 3 assumes 3 vCPUs executing on physical CPUs. Suppose vCPU 0 frequently demands data in GVP 3, which maps to GPP 8 and SPP 5, and that SPP 5 resides in off-chip DRAM. The hypervisor may want to migrate SPP 5 to die-stacked memory (e.g., SPP 512) to improve performance. On a VM exit (assumed to have occurred prior in time to Figure 3), the hypervisor modifies the nested page table to update the SPP, triggering translation coherence. There are three problems with this:

All vCPUs are identified as targets: Figure 3 shows that the hypervisor initiates translation coherence by setting the TLB flush request bit in every vCPU’s kvm_vcpu structure. kvm_vcpu stores vCPU state; when a vCPU is scheduled on a physical CPU, it provides register content, instruction pointers, etc. By setting these bits, the hypervisor signals that TLB, MMU cache, and nTLB entries need to be flushed.

Ideally, we would like the hypervisor to identify only the CPUs that actually cache the stale translation as targets. The hypervisor does spare physical CPUs that never executed the VM. However, it flushes all physical CPUs that ran any of the vCPUs of the VM, regardless of whether they cache the modified page table entries.

All vCPUs suffer VM exits: In the next step, the hypervisor launches inter-processor interrupts (IPIs) to all the vCPUs. IPIs use the processor’s advanced programmable interrupt controllers (APICs). APIC implementations vary; depending on the APIC technology, KVM converts broadcast IPIs into a loop of individual IPIs, or a loop across processor clusters. We have profiled the overheads of IPIs using microbenchmarks on Haswell systems, and like past work [10, 33], find that they are expensive, consuming thousands of clock cycles. If the receiving CPUs are running vCPUs, they suffer VM exits, compromising 3⃝ from Sec. 1. Targets then acknowledge the initiator, which is paused waiting for all vCPUs to respond.

All translation structures are flushed: The next step is to invalidate stale mappings in translation structure entries. Current architectures provide ISA and microarchitectural support for this via, for example, invlpg instructions in x86, etc. There are two caveats however. First, these instructions need the GVP of the modified nested page table mapping to identify the TLB entries that need to be invalidated. This is largely because modern TLBs maintain GVP bits in the tag. While this is a good design choice for non-virtualized systems, it is problematic for virtualized systems because hypervisors do not have easy access to GVPs. Instead, they have GPPs and SPPs. Consequently, KVM, Xen, etc., flush all TLB contents when they modify a nested page table entry, rather than selectively invalidating TLB entries. Second, there are currently no instructions to selectively invalidate MMU caches or nTLBs, even though they are tagged with GPPs and SPPs. The is because the marginal benefits of adding ISA support for selective MMU cache and nTLB invalidation are limited when the more performance-critical TLBs are flushed.

3.3 Hardware Versus Software Solutions

It is natural to ask whether translation coherence problems can be solved with smarter software. We have studied this possibility and have concluded that hardware solutions are superior. Fundamentally, software solutions only partially solve the problem of flushing all translation structures, and cannot solve the problem of identifying all vCPUs as translation coherence targets and prompting VM exits.

Consider the problem of flushing all translation structures. One might consider tackling this problem by modifying the guest-hypervisor interface to enable the hypervisor to use existing ISA support (e.g., invlpg instructions) to selectively invalidate TLB entries. But this only fixes TLB invalidation – no architectures today maintain selective invalidation instructions for MMU caches and nTLBs, so these would still have to be flushed.

Even if this problem could be solved, making target-side translation coherence handling lightweight is challenging. Fundamentally, handling translation coherence in software means that a context switch of the CPUs is unavoidable. One alternative to expensive VM exits might be to switch to lighterweight interrupts to query the guest OS for GVP-SPP mappings. Unfortunately, even these interrupts remain expensive. Specifically, we profiled interrupt costs using microbenchmarks on Intel’s Haswell machines and found that they require 640 cycles on average, which is just half of the average of 1300 cycles required for a VM exit. HATRIC, however, entirely eliminates these costs by never disrupting the operation of the guest OS or requiring context switching.

4 Hardware Design

We now detail HATRIC’s design, focusing mostly on hypervisor-initiated paging which modifies the nested page table. HATRIC achieves all three goals set out in Sec. 1. It does so by adding co-tags to translation structures to achieve precise invalidation. It then exposes these co-tags to the cache coherence protocol to precisely identify coherence targets and to eliminate VM exits.

4.1 Co-Tags

Figure 4: We add co-tags to store the system physical addresses where nested page table entries are stored. In our final implementation, we only store a subset of the system physical address bits.

We describe co-tags by discussing what they are, what they accomplish, how they are designed, and who sets them.

What are co-tags? Consider the page tables of Figure 4 and suppose that the hypervisor modifies the GPP 2-SPP 2 nested page table mapping, making the TLB entry caching information about SPP 2 stale. Since the TLB caches GVP-SPP mappings rather than GPP-SPP mappings, this means that we’d like to selectively invalidate GVP 1-SPP 2 from the TLB, and although not shown, corresponding MMU cache and nTLB entries. Co-tags allow us to do this by acting as tag extensions that allow precise identification of translations when the hypervisor does not know the GVP. Co-tags store the system physical address of the nested page table entry (nL1 from the bottom-most row in Figure 1). For example, GVP 1-SPP 2 uses the nested page table entry at system physical address 0x100c, which is stored in the co-tag.

What do co-tags accomplish? Co-tags not only permit precise translation information identification but can also be piggybacked on existing cache coherence protocols. When the hypervisor modifies a nested page table translation, cache coherence protocols detect the modification to the system physical address of the page table entry. Ordinarily, all private caches respond so that only one amongst them holds the up-to-date copy of the cache line storing the nested page table entry. With co-tags, HATRIC extends cache coherence as follows. Coherence messages, previously restricted to just private caches, are now also relayed to translation structures. Co-tags are used to identify which (if any) TLB, MMU cache, and nTLB entries correspond to the modified nested page table cache line. Overall, this means that co-tags: a⃝ pick up on nested page table changes entirely in hardware, without the need for IPIs, VM exits, or invlpg instructions; b⃝ rely on, without fundamentally changing, existing cache coherence protocols; c⃝ permit selective TLBs, MMU caches, and nTLBs rather than flushes.

How are co-tags implemented? Co-tags have one important drawback. System physical addresses on 64-bit systems require 8 bytes. If all 8 bytes are realized in the co-tag, each TLB entry doubles in size. MMU cache and nTLB entries triple in size. Since address translation can account for 13-15% of processor energy [56, 57, 58, 59, 60], these area and associated energy overheads are unacceptable.

Therefore, we decrease the resolution of co-tags, using fewer bits. This means that groups, rather than individual TLB entries may be invalidated when one nested page table entry is changed. However, judiciously-sized co-tags generally achieve a good balance between invalidation precision, and area/energy overheads. Sec. 6 shows, using detailed RTL modeling, that 2-byte co-tags (a per-core area overhead of 2%) strike a good balance. We specify the exact subset of address bits make up the co-tag in subsequent sections.

Who sets co-tags? For good performance, co-tags must be set by hardware without an OS or hypervisor interrupt. HATRIC uses the page table walker to do this. On TLB, MMU cache, and nTLB misses, the page table walker performs a two-dimensional page table walk. In so doing, it infers the system physical address of the page table entries and stores it in the TLB, MMU cache, and nTLB co-tags.

4.2 Integration with Cache Coherence

Modern cache coherence protocols can integrate not only readable and writable private caches, but also read-only instruction caches (though instruction caches do not have to be read-only). Since TLBs, MMU caches, and nested TLBs are fundamentally read-only structures, HATRIC integrates them into the existing cache coherence protocol in a manner similar to read-only instruction caches. Beyond this, HATRIC has minimal impact on the cache coherence protocol. We describe HATRIC’s operation on a directory-based MESI protocol, with the coherence directories located at the shared LLC cache banks. Without loss of generality, we use dual-grain coherence directories from recent work [61].

Translation structure coherence states: Since translation structures are read-only, their entries require only two coherence states: Shared (S), and Invalid (I). These two states may be realized using per-entry valid bits. When a translation is entered into the TLB, MMU cache, or nTLB, the valid bit is set, representing the S state; the translation can be accessed by the local CPU. The translation structure entry remains in this state until it receives a coherence message. Co-tags are compared to incoming messages; when an invalidation request matches the co-tag, the translation entry is invalidated.

Translation coherence initiators: Consider Figure 5. Before detailing the numbered transactions, let us consider HATRIC’s components. We show a 4-CPU system, with private L1 caches, 4 shared LLC banks, and per-bank coherence directories. We show TLBs and though they also exist, we omit MMU caches and nTLBs to save space. MMU caches and nTLBs interact with the cache coherence protocol in a manner that mirrors TLBs. We show 8 cached page table entries, represented as green and black boxes. Translation coherence is initiated by the hardware page table walker or OS/hypervisor software.

Figure 5: Coherence directories identify translation structures caching page table entries, aside from private L1 cache contents.

Page table walkers: These are hardware finite state machines that are invoked on TLB misses. Walkers traverse the page tables and are responsible for filling translation information into the translation structures and setting the co-tags. Walkers cannot map or unmap pages.

OS and hypervisor: These can traverse, map, and unmap page table entries using standard load/store instructions. HATRIC picks up these changes, and keeps all private cache and translation structures coherent.

Coherence directory: HATRIC minimally changes the coherence directory. Key design considerations are:

Directory entry changes: Figure 5 shows that the coherence directory tracks non-page table and page table cache lines. We make a minor change to directory entries, adding two bits to record whether cache lines belong to a guest page table (gPT) or nested page table (nPT). HATRIC uses these bits to identify the case when a line holding page table data is modified in the private caches. When this happens, coherence transactions need to be sent to the translation structures.

The nPT and gPT bits are set by the hardware page table walkers on fills to the TLBs, MMU caches, and nTLBs. One might initially expect this to be problematic in the case where the OS or hypervisor reads or writes a page table cache line in software. In reality however, this does not present correctness issues. Two situations are possible. In the first situation, the page table walker has previously accessed the cache line, and has already set the nPT or gPT bit in the cache line’s directory entry. There are no correctness issues in this case. In the second situation, the OS or hypervisor reads or writes a page table cache line that has previously never been looked up by the page table walker. In this case, there is actually no need to set the nPT or gPT bits in the coherence directory entry yet since no translations from this line are cached in the TLB, MMU cache, or nTLB anyway. Modifying the cache line at this point does not require coherence messages to be sent to the translation structures. When the page table walker does eventually access a translation from this cache line and fills it into the translation structures, it checks the access bit already maintained by x86-64 translation entries. The access bit records whether an entry has previously been filled into the TLB or accessed by the page table walker [62]. If this bit is clear, this means that the entry (and hence the cache line it resides in) has not been accessed by the page table walker yet. In this case, the page table walker sends a message to the coherence directory to update the nPT and gPT bits of the relevant cache line.

Coherence granularity: Figure 5 shows that directory entries store information at the cache line granularity. x86-64 systems cache 8 page table entries per 64-byte cache line. Hence, similar to false sharing in caches [63], HATRIC conservatively invalidates all translation structure entries caching these 8 page table entries, even if only a single page table entry is modified. For example, consider CPU 3 in Figure 5, where the TLB caches two translations mapped to the same cache line. If any CPU modifies either one of these translations, HATRIC has to invalidate both TLB entries. This has implications on the size of co-tags. Recall that in Sec. 4.1, we stated that co-tags use a subset of the address bits. We want use the least significant, and hence, highest entropy bits as co-tags. But since cache coherence protocols track groups of 8 translations, co-tags do not store the 3 least significant address bits. Our 2 byte co-tags use bits 19-3 of the system physical address storing the page table. Naturally, this means that translations from different addresses in the page table may alias to the same co-tag. In practice, this has little adverse affect on HATRIC’s performance.

Coherence specificity issues: To simplify hardware, coherence directories do not track where among the private caches, TLB, MMU cache, and nTLB the page table entries are cached. Instead, coherence directories are pseudo-specific. For example, Figure 5 shows that CPU 0 caches page table entries in the TLB and L1 cache, CPU 1 only caches them in the L1 cache, while CPU 3 only caches them in the TLB. Nevertheless, the coherence directory’s sharer list does not capture this distinction. Therefore, when a CPU modifies page table contents and invalidation messages need to be sent to the sharers, they are relayed to the L1 caches and all translation structures, regardless of which ones actually cache page tables. This results in spurious coherence activity (e.g., CPU 3’s L1 cache need not be relayed an invalidation message for any of the page table entries shown). In practice though, because modifications of the page table are rare compared to other coherence activity, this additional traffic is tolerable. Ultimately, the gains from eliminating high-latency software TLB coherence far outweigh these relatively minor overheads (see Sec. 6).

Cache and translation structure evictions: Directories track translations in a coarse-grained and pseudo-specific manner. This has important implications on cache line evictions. Ordinarily, when a private cache line is evicted, the coherence directory is relayed a message to update the line’s sharer list [61]. An up-to-date sharer list eliminates spurious coherence traffic to this line in the future. We continue to employ this strategy for non-page table cache lines but use a slightly different approach for page tables. When a cache line holding page table entries is evicted, its content may still be cached in the TLB, MMU cache, nTLB. Even worse, other translations with matching co-tags may still be residing in the translation structures. One option may be to detect all translations with matching co-tags and invalidate them. This hurts energy because of the additional translation structure lookups, and performance because of unnecessary TLB, MMU cache, and nTLB entry invalidations.

Figure 6: Coherence activity from the eviction of a cache line holding page table entries from CPU0’s private cache. HATRIC updates sharer list information lazily in response to cache line evictions.

Figure 6 shows how HATRIC handles this problem, contrasting it with traditional cache coherence. Suppose CPU 0 evicts a cache line with page table entries. Both approaches relay a message to the coherence directory. Ordinarily, we remove CPU 0 from the sharer list. However, if HATRIC sees that this message corresponds to a cache line storing a page table (by checking the directory entry’s page table bits), the sharer list is untouched. This means that if CPU 1 subsequently writes to the same cache line, HATRIC sends spurious invalidate messages to CPU 0, unlike traditional cache coherence. However, we mitigate frequency of spurious messages; when CPU 0 sees spurious coherence traffic, it sends a message back to the directory to demote CPU 0 from the sharer list. Sharer lists are hence lazily updated. For similar reasons, evictions from translation structures also lazily update coherence directory sharer lists.

Directory evictions: Finally, past work shows that coherence directory entry evictions require back-invalidations of the associated cache lines in the cores [61]. This is necessary for correctness; all lines in private caches must always have a directory entry. HATRIC extends this approach to relay back-invalidations to the TLBs, MMU caches, and nTLBs too.

4.3 Putting It All Together

Figure 5 details HATRIC’s overall operation. Initially, CPU 0’s TLB and L1 caches are empty. On a memory access, CPU 0 misses in the TLB and walks the page table 1⃝. Whenever a request is satisfied from a page table line in the L1 cache in the M, E, or S state, there is no need to initiate coherence transactions. However, suppose that the last memory reference in the page table walk from Figure 1 is absent in the L1 cache. A read request is sent to the coherence directory in step 2⃝.

Two scenarios are possible. In the first, the translation may be uncached in the private caches, and there is no coherence directory entry. A directory entry is allocated and the gPT or nPT bit is set. In the second scenario (shown in Figure 5), the request matches an existing directory entry. The nPT bit already is set and HATRIC reads the sharer list which identifies CPUs 1 and 3 as also caching the desired translation (and the 7 adjacent translations in the cache line) in shared state. In response, the cache line with the desired translations is sent back to CPU 0 (from CPU 1, 3, or memory, whichever is quicker), updating the L1 cache

3a
and TLB

3b
. Subsequently, the sharer list adds CPU 0.

Now suppose that CPU 1 runs the hypervisor and unmaps the solid green translation from the nested page table in step 4⃝. To transition the L1 cache line into the M state, the cache coherence protocol relays a message to the coherence directory. The corresponding directory entry is identified in 5⃝, and we find that CPU 0 and 3 need to be sent invalidation requests. However, the sharer list is (i) coarse-grained and (ii) pseudo-specific. Because of (i), CPU 0 has to invalidate not only its TLB entry

6a
but also 8 translations in the L1 cache

6b
, and CPU 3 has to invalidate the 2 TLB entries with matching co-tags

6c
. Because of (ii), CPU 1’s L1 cache receives a spurious invalidation message

6d
.

4.4 Other Key Observations

Scope: HATRIC is applicable to virtualized and non-virtualized systems. For the latter, the co-tags may simply be used to store the physical addresses of page tables. Further, while we have focused on nested page table coherence, HATRIC can also be trivially modified to support shadow page tables too [40]. The co-tags merely have to store the memory addresses where shadow page tables are stored.

Metadata updates: Beyond software changes to the translations, they may also be changed by hardware page table walkers. Specifically, page table walkers update dirty and access bits to aid page replacement policies [22]. But since these updates are picked up by the standard cache coherence protocol, HATRIC naturally handles these updates too.

Prefetching optimizations: Beyond simply invalidating stale translation structure entries, HATRIC could potentially directly update (or prefetch) the updated mappings into the translation structures. Since a thorough treatment of these studies requires an understanding of how to manage translation access bits while speculatively prefetching into translation structures [62], we leave this for future work.

Coherence protocols: We have studied a MESI directory based coherence protocol but we have also implemented HATRIC atop MOESI protocols too, as well as snooping protocols like MESIF [64]. HATRIC requires no fundamental changes to support these protocols.

Synonyms and superpages: HATRIC naturally handles synonyms or virtual address aliases. This is because synonyms are defined by unique translations in separate page table locations, and hence separate system physical addresses. Therefore, changing or removing a translation has no impact on other translations in the synonym set, allowing HATRIC to be agnostic to synonyms. Similarly, HATRIC supports superpages, which also occupy unique translation entries and can hence be easily detected by co-tags.

Multiprogrammed workloads: One might expect that when an application’s physical page is remapped, there is no need for translation coherence activities to the other applications, because they operate on distinct address spaces. Unfortunately, however, hypervisors do not know which physical CPUs an application executed on; all they know is the vCPUs and the physical the entire VM uses. Therefore, the hypervisor conservatively flushes the even the translation structures of CPUs that never ran the offending application. HATRIC completely eliminates this problem by precisely tracking the correspondence between translations and CPUs.

Comparison to past approaches: HATRIC is inspired by past work on UNITD [32]. Like HATRIC, UNITD piggybacks translation coherence atop cache coherence protocols. Unlike HATRIC however, UNITD cannot support virtualized systems or MMU cache and nTLB coherence. Further, HATRIC uses energy-frugal co-tags instead of UNITD’s large reverse-lookup CAM circuitry, achieving far greater energy efficiency. We showcase this in Sec. 6 where we compare the efficiency of HATRIC versus an enhanced UNITD design for virtualization. Beyond UNITD, past work on DiDi [33] also targets translation coherence for non-virtualized systems. Similarly, recent work investigates translation coherence overheads in the context of die-stacked DRAM [10]. While this work mitigates translation coherence overheads, it does so specifically for non-virtualized x86 architectures, and ignores MMU caches and nTLBs. Finally, recent work uses software mechanisms to reduce translation overheads for guest page table modifications [47], while HATRIC also solves the problem of nested page table coherence.

5 Methodology

Our experimental methodology has two steps. First, we modify KVM to implement paging on a two-level memory with die-stacked DRAM. Second, we use detailed cycle-accurate simulation to assess performance and energy.

5.1 Die-Stacked DRAM Simulation

We evaluate HATRIC’s performance on a detailed cycle-accurate simulation framework that models the operation of a 32-CPU Haswell processor. We assume 2GB of die-stacked DRAM with 4 the bandwidth of slower 8GB off-chip DRAM, similar to prior work [10]. Each CPU maintains 32KB L1 caches, 256KB L2 caches, 64-entry L1 TLBs, 512-entry L2 TLBs, 32-entry nTLBs [38], and 48-entry paging structure MMU caches [37]. Further, we assume a 20MB LLC. We model the energy usage of this system using the CACTI framework [65]. We use Ubuntu 15.10 Linux as our guest OS. Further, we evaluate HATRIC in detail using KVM. Beyond this, we have also run Xen to highlight HATRIC’s generality with other hypervisors.

We use a trace-based approach to drive our simulation framework. We collect instruction traces from our modified hypervisors with 50 billion memory references using a modified version of Pin which tracks all GVPs, GPPs, and SPPs, as well as changes to the guest and nested page tables. In order to collect accurate paging activity, we collect these traces on a real-system. Ideally, we would like this system to use die-stacked DRAM but since this technology is in its infancy, we are inspired by recent work [10] to modify a real-system to mimic the activity of die-stacking. We take an existing multi-socket NUMA platform, and by introducing contention, creates two different speeds of DRAM. We use a 2-socket Intel Xeon E5-2450 system, running our software stack. We dedicate the first socket for execution of the software stack and mimicry of fast or die-stacked DRAM. The second socket mimics the slow or off-chip DRAM. It does so by running several instances of memhog on its cores. Similar to prior work [28, 22, 43], we use memhog to carefully generate memory contention to achieve the desired bandwidth differential between the fast and slow DRAM of 4. By using Pin to track KVM and Linux paging code on this infrastructure, we accurately generate instruction traces to test HATRIC.

5.2 KVM Paging Policies

Our goal is to showcase the overheads imposed by translation coherence on paging decisions rather than design the optimal paging policy, leaving this for future work. So, we pick well-known paging policies that cover a wide range of design options. For example, we have studied FIFO and LRU replacement policies, finding the latter to perform better, as expected. We implement LRU policies in KVM by repurposing Linux’s well-known pseudo-LRU CLOCK policy [66]. LRU alone doesn’t always provide good performance since it is expensive to traverse page lists to identify good candidates for eviction from die-stacked memory. Instead, performance is improved by moving this operation off the critical path of execution; we therefor pre-emptively evict pages from die-stacked memory so that a pool of free pages are always maintained. We call this migration daemon and combine it with LRU. We have also investigated the benefits of page prefetching; that is, when an application demand fetches a page from off-chip to die-stacked memory, we also prefetch a set number of adjacent pages. Generally, we have found that the best paging policy uses a combination of these approaches.

5.3 Workloads

Our focus is on two sets of workloads. The first set comprises applications that benefit from the higher bandwidth of die-stacked memory. We use canneal and facesim from the Parsec suite [67], data caching and tunkrank from Cloudsuite [68], and graph500 as part of this group. We also create 80 multiprogrammed combinations of workloads from all the Spec applications to showcase the problem of imprecise target identification in virtualized translation coherence.

Our second group of workloads is made up of smaller-footprint applications whose data largely fits within the die-stacked DRAM. We use these workloads to evaluate HATRIC’s overheads in situations where hypervisor-mediated paging (and hence translation coherence) between die-stacked and off-chip DRAM is rarer. We use the remaining Parsec applications, and Spec applications for these studies.

6 Evaluation

Figure 7: For varying vCPUs, runtime of the best KVM paging policy without HATRIC (sw), with HATRIC (hatric), and with zero-overhead translation coherence (ideal). All results are normalized to the case without die-stacked DRAM.

Performance as a function of vCPU counts: Figure 7 shows HATRIC’s runtime, normalized as a fraction of application runtime in the absence of any die-stacked memory (no-hbm from Figure 2). We compare runtimes for the best KVM paging policies (sw), HATRIC, and ideal unachievable zero-overhead translation coherence (ideal). Further, we vary the number of vCPUs per VM and observe the following.

HATRIC is always within 2-4% of the ideal performance. In some cases, HATRIC is instrumental in achieving any gains from die-stacked memory at all. Consider data caching, which slows down when using die-stacked memory, because of translation coherence overheads. HATRIC cuts runtimes down to roughly 75% of the baseline runtime in all cases.

Figure 7 also shows that HATRIC is valuable at all vCPU counts. In some cases, more vCPUs exacerbate translation coherence overheads. This is because IPI broadcasts become more expensive and more vCPUs suffer VM exits. This is why, for example, data caching and tunkrank become slower (see sw) when vCPUs increase from 4 to 8. HATRIC eliminates these problems, flattening runtime improvements across vCPU counts. In other scenarios, fewer vCPUs worsen performance since each vCPU performs more of the application’s total work. Here, the impact of a full TLB, nTLB, and MMU cache flush for every page remapping is very expensive (e.g., graph500 and facesim). Here, HATRIC again eliminates these overheads almost entirely.

Figure 8: HATRIC’s performance benefits as a function of KVM paging policies, with LRU, migration daemons (mig-dmn), and prefetching (pref.). Results are normalized to the case without die-stacked DRAM.

Performance as a function of paging policy: Figure 8 also shows HATRIC performance, but this time as a function of different KVM paging policies. We study three policies with 16 vCPUs. First, we show lru, which determines which pages to evict from die-stacked DRAM. We then add the migration daemon (&mig-dmn), and page prefetching (&pref).

Figure 8 shows HATRIC improves runtime substantially for any paging policy. Performance is best when all techniques are combined, but HATRIC achieves 10-30% performance improvements even for just lru. Furthermore, Figure 8 shows that translation coherence overheads can often be so high that the paging policy itself makes little difference to performance. Consider tunkrank, where the difference between lru versus the &pref bars is barely 2-3%. With HATRIC, however, paging optimizations like prefetching and migration daemons help.

Figure 9: HATRIC’s performance benefits as a function of translation structure size. 1 indicates default sizes, 2 doubles sizes, and so on. All results are normalized to the case without die-stacked DRAM.

Impact of translation structure sizes: One of HATRIC’s advantages is that it converts translation structure flushes to selective invalidations. This improves TLB, MMU cache, and nTLB hit rates substantially, obviating the need for expensive two-dimensional page table walks. We expect HATRIC to improve performance even more as translation structures become bigger (and flushes needlessly evict more entries). Figure 9 quantifies the relationship. We vary TLB, nTLB, and MMU cache sizes from the default (see Sec. 5) to double (2) and quadruple (4) the number of entries.

Figure 9 shows that translation structure flushes largely counteract the benefits of greater size. Specifically, the sw results see barely any improvement, even when sizes are quadrupled. Inter-DRAM page migrations essentially flush the translation structures so often that additional entries are not effectively leveraged. Figure 9 shows that this is a wasted opportunity since zero-overhead translation coherence (ideal) actually does enjoy 5-7% performance benefits. HATRIC solves this problem, comprehensively achieving within 1% of the ideal, thereby exploiting larger translation structures.

Figure 10: (Left) Weighted runtime for all 80 multi-programmed workloads on VMs without (actual) and with HATRIC (hatric); (right) the same for the slowest application in mix.

Multi-programmed workloads: We now focus on multiprogrammed workloads made up sequential applications. Each workload runs 16 Spec benchmarks on a Linux VM atop KVM. As is standard for multiprogrammed workloads, we use two performance metrics [69, 70]. The first is weighted runtime improvement, which captures overall system performance. The second is the runtime improvement of the slowest application in the workload, capturing fairness.

Figure 10 shows our results. The graph on the left plots the weighted runtime improvement, normalized to cases without die-stacked DRAM. As usual, sw represents the best KVM paging policy. The x-axis represents the workloads, arranged in ascending order of runtime. The lower the runtime, the better the performance. Similarly, the graph on the right of Figure 9 shows shows the runtime of the slowest application in the workload mix; again, lower runtimes indicate a speedup in the slowest application.

Figure 10 shows that translation coherence can be disastrous to the performance of multiprogrammed workloads. More than 70% of the workload combinations suffer performance degradation with die-stacking. These applications suffer from unnecessary translation structure flushes and VM exits, caused by software translation coherence’s imprecise target identification. Runtime is more than 2 for 11 workloads. Additionally, translation coherence degrades application fairness. For example, in more than half the workloads, the slowest application’s runtime is (2)+ with a maximum of (4)+. Applications that struggle are usually those with limited memory-level parallelism that benefit little from the higher bandwidth of die-stacked memory and instead, suffer from the additional translation coherence overheads.

HATRIC solves all these issues, achieving improvements for every single weighted runtime, and even for each of the slowest applications. In fact, HATRIC entirely eliminating translation coherence overheads, reducing runtime to 50-80% of the baseline without die-stacked DRAM. The key enabler is HATRIC’s precise identification of coherence targets – applications that do not need to participate in translation coherence operations have their translation structure contents left unflushed and do not suffer VM exits.

Figure 11: (Left) Performance-energy plots for default HATRIC configuration compared to a baseline with the best paging policy; and (right) impact of co-tag size on performance-energy tradeoffs.

Performance-energy tradeoffs: Intuitively, we expect that since HATRIC reduces runtime substantially, it should reduce static energy sufficiently to offset the higher energy consumption from the introduction of co-tags. Indeed, this is true for workloads that have sufficiently large memory footprints to trigger inter-memory paging. However, we also assess HATRIC’s energy implications on workloads that do not frequently remap pages (i.e., their memory footprints fit comfortably within die-stacked DRAM).

The graph on the left of Figure 11 plots all the workloads including the single-threaded and multithreaded ones that benefit from die-stacking and those whose memory needs fit entirely in die-stacked DRAM. The x-axis plots the workload runtime, as a fraction of the runtime of sw results. The y-axis plots energy, similarly normalized. We desire points that converge towards the lower-left corner of the graph.

The graph on the left of Figure 11 shows that HATRIC always boosts performance, and almost always improves energy too. Energy savings of 1-10% are routine. In fact, HATRIC even improves the performance and energy of many workloads that do not page between the two memory levels. This is because these workloads still remap pages to defragment memory (to support superpages) and HATRIC mitigates the associated translation coherence overheads. There are some rare instances (highlighted in black) where energy does exceed the baseline by 1-1.5%. These are workloads for whom efficient translation coherence does not make up for the additional energy of the co-tags. Nevertheless, these overheads are low, and their instances rare.

Co-tag sizing: We now turn to co-tag sizing. Excessively large co-tags consume significant lookup and static energy, while small ones force HATRIC to invalidate too many translation structures on a page remap. The graph on the right of Figure 11 shows the performance-energy implications of varying co-tag size from 1 to 3 bytes.

First and foremost, 2B co-tags – our design choice – provides the best balance of performance and energy. While 3B co-tags track page table entries at a finer granularity, they only modestly improve performance over 2B co-tags, but consume much more energy. Meanwhile 1B co-tags suffer in terms of both performance and energy. Since 1B co-tags have a coarser tracking granularity, they invalidate more translation entries from TLBs, MMU caches, and nTLBs than larger co-tags. And while the smaller co-tags do consume less lookup and static energy, these additional invalidations lead to more expensive two-dimensional page table walks and a longer system runtime. The end result is an increase in energy too.

Coherence directory design decisions: Sec. 4 detailed the nuances modifying traditional coherence directories to support translation coherence. Figure 12 captures the performance and energy (normalized to those of the best paging policy or sw in previous graphs) of these approaches. We consider the following options, beyond baseline HATRIC.

Figure 12: Baseline HATRIC versus approaches with eager update of directory on cache and translation structure evictions (EGR-dir-update), fine-grained tracking of translations (FG-tracking), and an infinite directory with no back-invalidations (No-back-inv). All comines these approach. We show average runtime and energy, normalized to the metrics for the best paging policy without HATRIC.

EGR-dir-update: This is a design that eagerly updates coherence directories whenever a translation entry is evicted from a CPU’s L1 cache or translation structures. While this does reduce spurious coherence messages, it requires expensive lookups in translation structures to ensure that entries with the same co-tag have been evicted. Figure 12 shows that the performance gains from reduced coherence traffic is almost negligible, while energy does increase, relative to HATRIC.

FG-tracking: We study a hypothetical design with greater specificity in translation tracking. That is, coherence directories are modified to track whether translations are cached in the TLBs, MMU caches, nTLBs, or L1 caches. Unlike HATRIC, if a translation is cached only in the MMU cache but not the TLB, the latter is not sent invalidation requests. Figure 12 shows that while one might expect this specificity to result in reduced coherence traffic, system energy is actually slightly higher than HATRIC. This is because more specificity requires more complex and area/energy intensive coherence directories. Further, since the runtime benefits are small, we believe HATRIC remains the smarter choice.

No-back-inv: We study an unrealistically ideal design with infinitely-sized coherence directories which never need to relay back-invalidations to private caches or translation structures. We find that this does reduce energy and runtime, but not significantly from HATRIC’s dual-grain coherence directory based on [61].

All: Figure 12 compares HATRIC to an approach which marries all the optimizations discussed. HATRIC almost exactly meets the same performance and is actually more energy-efficient, largely because the eager updates of coherence directories add significant translation structure lookup energy.

Comparison with UNITD: We now compare HATRIC to prior work on UNITD [32]. To do this, we first upgrade the baseline UNITD design in several ways. First, and most importantly, we extend add support for virtualization by storing the system physical address of nested page tables entries are stored in the reverse-lookup CAM originally proposed [32]. Second, we extend UNITD to work seamlessly with coherence directories. We call this upgraded design UNITD++.

Figure 13 compares HATRIC and UNITD++ results, normalized to results from the case without die-stacked DRAM. As expected, both approaches outperform a system with only traditional software-based translation coherence (sw). However, HATRIC typically provides an additional 5-10% performance boost versus UNITD++ by also extending the benefits of hardware translation coherence to MMU caches and nTLBs. Further, HATRIC is more energy efficient than UNITD++ as it boosts performance (saving static energy) but also does not need reverse-lookup CAMs.

Figure 13: Comparison of HATRIC’s performance and energy versus UNITD++. All results are normalized to results for a system without die-stacked memory and compared to sw.

Xen results: In order to assess HATRIC’s generality across hypervisors, we have begun studying it’s effectiveness on Xen. Because our memory traces require months to collect, we have thus far evaluated canneal and data caching, assuming 16 vCPUs. Our initial results show that Xen’s performance is improved by 21% and 33% for canneal and data caching respectively, over the best paging policy employing software translation.

7 Conclusion

We present a case for folding translation coherence atop existing hardware cache coherence protocols. We achieve this with simple modifications to translation structures (TLBs, MMU caches, and nTLBs) and with state-of-the-art coherence protocols. Our solutions are general (they support nested and guest page table modifications) and readily-implementable. We believe, therefore, that HATRIC will become essential for upcoming systems, especially as they rely on page migration to exploit heterogeneous memory systems.

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