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37\title{On the correctness of an optimising assembler for the Intel MCS-51 microprocessor\thanks{The project CerCo acknowledges the financial support of the Future and Emerging Technologies (FET) programme within the Seventh Framework Programme for Research of the European Commission, under FET-Open grant number: 243881.}}
38\author{Dominic P. Mulligan \and Claudio Sacerdoti Coen}
39\institute{Dipartimento di Scienze dell'Informazione,\\ Universit\'a degli Studi di Bologna}
40
41\bibliographystyle{splncs03}
42
43\begin{document}
44
45\maketitle
46
47\begin{abstract}
48We present a proof of correctness in Matita for an optimising assembler for the MCS-51 microcontroller.
49The efficient expansion of pseudoinstructions, namely jumps, into machine instructions is complex.
50We isolate the decision making over how jumps should be expanded from the expansion process itself as much as possible using `policies', making the proof of correctness for the assembler more straightforward.
51
52%We observe that it is impossible for an optimising assembler to preserve the semantics of every assembly program.
53%Assembly language programs can manipulate concrete addresses in arbitrary ways.
54Our proof strategy contains a tracking facility for `good addresses' and only programs that use good addresses have their semantics preserved under assembly, as we observe that it is impossible for an assembler to preserve the semantics of every assembly program.
55Our strategy offers increased flexibility over the traditional approach to proving the correctness of assemblers, wherein addresses in assembly are kept opaque and immutable.
56In particular, we may experiment with allowing the benign manipulation of addresses.
57\keywords{Verified software, CerCo (Certified Complexity), MCS-51 microcontroller, Matita proof assistant}
58\end{abstract}
59
60% ---------------------------------------------------------------------------- %
61% SECTION                                                                      %
62% ---------------------------------------------------------------------------- %
63\section{Introduction}
64\label{sect.introduction}
65
66We consider the formalisation of an assembler for the Intel MCS-51 8-bit microprocessor in the Matita proof assistant~\cite{asperti:user:2007}.
67This formalisation forms a major component of the EU-funded CerCo (`Certified Complexity') project~\cite{cerco:2011}, concerning the construction and formalisation of a concrete complexity preserving compiler for a large subset of the C programming language.
68
69The MCS-51 dates from the early 1980s and is commonly called the 8051/8052.
70Derivatives are still widely manufactured by a number of semiconductor foundries, with the processor being used especially in embedded systems.
71
72The MCS-51 has a relative paucity of features compared to its more modern brethren, with the lack of any caching or pipelining features meaning that timing of execution is predictable, making the MCS-51 very attractive for CerCo's ends.
73However, the MCS-51's paucity of features---though an advantage in many respects---also quickly becomes a hindrance, as the MCS-51 features a relatively minuscule series of memory spaces by modern standards.
74As a result our C compiler, to be able to successfully compile realistic programs for embedded devices, ought to produce `tight' machine code.
75
76To do this, we must solve the `branch displacement' problem---deciding how best to expand pseudojumps to labels in assembly language to machine code jumps.
77The branch displacement problem arises when pseudojumps can be expanded
78in different ways to real machine instructions, but the different expansions
79are not equivalent (e.g. differ in size or speed) and not always
80correct (e.g. correctness is only up to global constraints over the compiled
81code). For instance, some jump instructions (short jumps) are very small
82and fast, but they can only reach destinations within a
83certain distance from the current instruction. When the destinations are
84too far away, larger and slower long jumps must be used. The use of a long jump may
85augment the distance between another pseudojump and its target, forcing
86another long jump use, in a cascade. The job of the optimising
87compiler (assembler) is to individually expand every pseudo-instruction in such a way
88that all global constraints are satisfied and that the compiled program
89is minimal in size and faster in concrete time complexity.
90This problem is known to be computationally hard for most CISC architectures (see~\cite{hyde:branch:2006}).
91
92To simplify the CerCo C compiler we have chosen to implement an optimising assembler whose input language the compiler will target.
93Labels, conditional jumps to labels, a program preamble containing global data and a \texttt{MOV} instruction for moving this global data into the MCS-51's one 16-bit register all feature in our assembly language.
94We further simplify by ignoring linking, assuming that all our assembly programs are pre-linked.
95
96Another complication we have addressed is that of the cost model.
97CerCo imposes a cost model on C programs or, more specifically, on simple blocks of instructions.
98This cost model is induced by the compilation process itself, and its non-compositional nature allows us to assign different costs to identical C statements depending on how they are compiled.
99In short, we aim to obtain a very precise costing for a program by embracing the compilation process, not ignoring it.
100At the assembler level, this is reflected by our need to induce a cost
101model on the assembly code as a function of the assembly program and the
102strategy used to solve the branch displacement problem. In particular, the
103optimising compiler should also return a map that assigns a cost (in clock
104cycles) to every instruction in the source program. We expect the induced cost
105to be preserved by the compiler: we will prove that the compiled code
106tightly simulates the source code by taking exactly the predicted amount of
107time.
108
109Note that the temporal tightness of the simulation is a fundamental prerequisite
110of the correctness of the simulation because some functions of the MCS-51---timers and I/O---depend on the microprocessor's clock.
111If the pseudo- and concrete clock differ the result of an I/O operation may not be preserved.
112
113Branch displacement algorithms must have a deep knowledge of the way
114the rest of the assembler works in order to build globally correct solutions.
115Proving their correctness is quite a complex task (see, for instance,
116the compaion paper~\cite{boender:correctness:2012}).
117Nevertheless, the correctness of the whole assembler only depends on the
118correctness of the branch displacement algorithm.
119Therefore, in the rest of the paper, we presuppose the
120existence of a correct policy, to be computed by a branch displacement
121algorithm if it exists. A policy is the decision over how
122any particular jump should be expanded; it is correct when the global
123constraints are satisfied.
124The assembler fails to assemble an assembly program if and only if a correct policy does not exist.
125This is stated in an elegant way in the dependent type of the assembler: the assembly function is total over a program, a policy and the proof that the policy is correct for that program.
126
127A final complication in the proof is due to the kind of semantics associated to pseudo-assembly programs.
128Should assembly programs be allowed to freely manipulate addresses?
129The traditional answer is `no': values stored in memory or registers are either
130concrete data or symbolic addresses. The latter can only be manipulated
131in very restricted ways and programs that do not do so are not assigned a semantics and cannot be reasoned about.
132All programs that have a semantics have it preserved by the assembler.
133We take an alternative approach, allowing programs to freely
134manipulate addresses non-symbolically but only granting a preservation of semantics
135to those programs that act in `well-behaved' ways.
136In principle, this should allow some reasoning on the actual semantics of malign programs.
137In practice, we note how our approach facilitates more code reuse between the semantics of assembly code and object code.
138
139The rest of this paper is a detailed description of our proof that is marginally still a work in progress.
140\paragraph{Matita}
141Matita is a proof assistant based on a variant of the Calculus of (Co)inductive Constructions~\cite{asperti:user:2007}.
142It features dependent types that we exploit in the formalisation.
143The (simplified) syntax of the statements and definitions in the paper should be self-explanatory.
144Pairs are denoted with angular brackets, $\langle-, -\rangle$.
145
146Matita features a liberal system of coercions.
147It is possible to define a uniform coercion $\lambda x.\langle x,?\rangle$ from every type $T$ to the dependent product $\Sigma x:T.P~x$.
148The coercion opens a proof obligation that asks the user to prove that $P$ holds for $x$.
149When a coercion must be applied to a complex term (a $\lambda$-abstraction, a local definition, or a case analysis), the system automatically propagates the coercion to the sub-terms
150 For instance, to apply a coercion to force $\lambda x.M : A \to B$ to have type $\forall x:A.\Sigma y:B.P~x~y$, the system looks for a coercion from $M: B$ to $\Sigma y:B.P~x~y$ in a context augmented with $x:A$.
151This is significant when the coercion opens a proof obligation, as the user will be presented with multiple, but simpler proof obligations in the correct context.
152In this way, Matita supports the `Russell' proof methodology developed by Sozeau in~\cite{sozeau:subset:2006}, with an implementation that is lighter and more tightly integrated with the system than that of Coq.
153
154% ---------------------------------------------------------------------------- %
155% SECTION                                                                      %
156% ---------------------------------------------------------------------------- %
157\section{The proof}
158\label{sect.the.proof}
159
160The aim of the section is to explain the main ideas and steps of the certified
161proof of correctness for an optimizing assembler for the MCS-51.
162
163In Section~\ref{subsect.machine.code.semantics} we sketch an operational semantics (a realistic and efficient emulator) for the MCS-51.
164We also introduce a syntax for decoded instructions that will be reused for the assembly language.
165
166In Section~\ref{subsect.assembly.code.semantics} we describe the assembly language and its operational semantics.
167The latter is parametric in the cost model that will be induced by the assembler, reusing the semantics of the machine code on all `real' instructions.
168
169Branch displacement policies are introduced in Section~\ref{subsect.the.assembler} where we also describe the assembler as a function over policies as previously described.
170
171To prove our assembler correct we show that the object code given in output, together with a cost model for the source program, simulates the source program executed using that cost model.
172The proof can be divided into two main lemmas.
173The first is correctness with respect to fetching, described in Section~\ref{subsect.total.correctness.of.the.assembler}.
174Roughly it states that a step of fetching at the assembly level, returning the decoded instruction $I$, is simulated by $n$ steps of fetching at the object level that returns instructions $J_1,\ldots,J_n$, where $J_1,\ldots,J_n$ is, amongst the possible expansions of $I$, the one picked by the policy.
175The second lemma states that $J_1,\ldots,J_n$ simulates $I$ but only if $I$ is well-behaved, i.e. manipulates addresses in `good' ways.
176To keep track of well-behaved address manipulations we record where addresses are currently stored (in memory or an accumulator).
177We introduce a dynamic checking function that inspects this map to determine if the operation is well-behaved, with an affirmative answer being the pre-condition of the lemma.
178The second lemma is detailed in Section~\ref{subsect.total.correctness.for.well.behaved.assembly.programs} where we also establish correctness of our assembler as a composition of the two lemmas: programs that are well-behaved when executed under the cost model induced by the compiler are correctly simulated by the compiled code.
179
180% ---------------------------------------------------------------------------- %
181% SECTION                                                                      %
182% ---------------------------------------------------------------------------- %
183
184\subsection{Machine code and its semantics}
185\label{subsect.machine.code.semantics}
186
187We implemented a realistic and efficient emulator for the MCS-51 microprocessor.
188An MCS-51 program is just a sequence of bytes stored in the read-only code
189memory of the processor, represented as a compact trie of bytes addressed
190by the program counter.
191The \texttt{Status} of the emulator is described as
192a record that contains the microprocessor's program counter, registers, stack
193pointer, clock, special function registers, code memory, and so on.
194The value of the code memory is a parameter of the record since it is not
195changed during execution.
196
197The \texttt{Status} records is itself an instance of a more general
198datatype \texttt{PreStatus} that abstracts over the implementation of code
199memory in order to reuse the same datatype for the semantics of the assembly
200language in the next section.
201
202The execution of a single instruction is performed by the \texttt{execute\_1}
203function, parametric over the content \texttt{cm} of the code memory:
204\begin{lstlisting}
205definition execute_1: $\forall$cm. Status cm $\rightarrow$ Status cm
206\end{lstlisting}
207
208The function \texttt{execute\_1} closely matches the fetch-decode-execute
209cycle of the MCS-51 hardware, as described by a Siemen's manufacturer's data sheet~\cite{siemens:2011}.
210Fetching and decoding are performed simultaneously:
211we first fetch, using the program counter, from code memory the first byte of the instruction to be executed, decoding the resulting opcode, fetching more bytes as is necessary to decode the arguments.
212Decoded instructions are represented by the \texttt{instruction} data type
213which extends a data type of \texttt{preinstruction}s that will be reused
214for the assembly language.
215\begin{lstlisting}
216inductive preinstruction (A: Type[0]): Type[0] :=
217 | ADD: $\llbracket$acc_a$\rrbracket$ → $\llbracket$registr; direct; indirect; data$\rrbracket$ $\rightarrow$ preinstruction A
218 | DEC: $\llbracket$acc_a; registr; direct; indirect$\rrbracket$ $\rightarrow$ preinstruction A
219 | JB: $\llbracket$bit_addr$\rrbracket$ $\rightarrow$ A $\rightarrow$ preinstruction A
220 | ...
221
222inductive instruction: Type[0] :=
223 | LCALL: $\llbracket$addr16$\rrbracket$ $\rightarrow$ instruction
224 | AJMP: $\llbracket$addr11$\rrbracket$ $\rightarrow$ instruction
225 | RealInstruction: preinstruction $\llbracket$relative$\rrbracket$ $\rightarrow$ instruction.
226 | ...
227\end{lstlisting}
228The MCS-51 has many operand modes, but an unorthogonal instruction set: every
229opcode is only enable for a finite subset of the possible operand modes.
230Here we exploit dependent types and an implicit coercion to synthesize
231the type of arguments of opcodes from a vector of names of operand modes.
232For example, \texttt{ACC} has two operands, the first one constrained to be
233the \texttt{A} accumulator, and the second one to be a disjoint union of
234register, direct, indirect and data operand modes.
235
236The parameterised type $A$ of \texttt{preinstruction} represents the addressing mode allowed for conditional jumps; in the \texttt{RealInstruction} constructor
237we constraint it to be a relative offset. A different instantiation will be
238used in the next Section for assembly programs.
239
240Once decoded, execution proceeds by a case analysis on the decoded instruction, following the operation of the hardware.
241For example, the \texttt{DEC} preinstruction (`decrement') is executed as follows:
242\begin{lstlisting}
243 | DEC addr $\Rightarrow$
244  let s := add_ticks1 s in
245  let $\langle$result, flags$\rangle$ := sub_8_with_carry (get_arg_8 s true addr)
246   (bitvector_of_nat 8 1) false in
247     set_arg_8 s addr result
248\end{lstlisting}
249
250Here, \texttt{add\_ticks1} models the incrementing of the internal clock of the microprocessor; it is a parameter of the semantics of \texttt{preinstruction}s
251that is fixed in the semantics of \texttt{instruction}s according to the
252manufacturer datasheet.
253
254% ---------------------------------------------------------------------------- %
255% SECTION                                                                      %
256% ---------------------------------------------------------------------------- %
257
258\subsection{Assembly code and its semantics}
259\label{subsect.assembly.code.semantics}
260
261An assembly program is a list of potentially labelled pseudoinstructions, bundled with a preamble consisting of a list of symbolic names for locations in data memory (i.e. global variables).
262All preinstructions are pseudoinstructions, but conditional jumps are now
263only allowed to use \texttt{Identifiers} (labels) as their target.
264\begin{lstlisting}
265inductive pseudo_instruction: Type[0] :=
266  | Instruction: preinstruction Identifier $\rightarrow$ pseudo_instruction
267    ...
268  | Jmp: Identifier $\rightarrow$ pseudo_instruction
269  | Call: Identifier $\rightarrow$ pseudo_instruction
270  | Mov: $\llbracket$dptr$\rrbracket$ $\rightarrow$ Identifier $\rightarrow$ pseudo_instruction.
271\end{lstlisting}
272The pseudoinstructions \texttt{Jmp}, \texttt{Call} and \texttt{Mov} are generalisations of machine code unconditional jumps, calls and move instructions respectively, all of whom act on labels, as opposed to concrete memory addresses.
273The object code calls and jumps that act on concrete memory addresses are ruled
274out of assembly programs not being included in the preinstructions (see previous
275Section).
276
277Execution of pseudoinstructions is an endofunction on \texttt{PseudoStatus}.
278A \texttt{PseudoStatus} is an instance of \texttt{PreStatus} that differs
279from a \texttt{Status} only in the datatype used for code memory: a list
280of optionally labelled pseudoinstructions versus a trie of bytes.
281The \texttt{PreStatus} type is crucial for sharing the majority of the
282semantics of the two languages.
283
284Emulation for pseudoinstructions is handled by \texttt{execute\_1\_pseudo\_instruction}:
285\begin{lstlisting}
286definition execute_1_pseudo_instruction:
287 $\forall$cm. $\forall$costing:($\forall$ppc: Word. ppc < $\mid$snd cm$\mid$ $\rightarrow$ nat $\times$ nat).
288  $\forall$s:PseudoStatus cm. program_counter s < $\mid$snd cm$\mid$ $\rightarrow$ PseudoStatus cm
289\end{lstlisting}
290The type of \texttt{execute\_1\_pseudo\_instruction} is more involved than
291that of \texttt{execute\_1}. The first difference is that execution is only
292defined when the program counter points to a valid instruction, i.e.
293it is smaller than the length $\mid$\texttt{snd cm}$\mid$ of the program.
294The second difference is the abstraction over the cost model, abbreviated
295here as \emph{costing}.
296The costing is a function that maps valid program counters to pairs of natural numbers representing the number of clock ticks used by the pseudoinstructions stored at those program counters. For conditional jumps the two numbers differ
297to represent different costs for the `true branch' and the `false branch'.
298In the next Section we will see how the optimizing
299assembler induces the only costing that is preserved by compilation.
300Obviously the induced costing is determined by the branch displacement policy
301that decides how to expand every pseudojump to a label into concrete
302instructions.
303
304Execution proceeds by first fetching from pseudo-code memory using the program counter---treated as an index into the pseudoinstruction list.
305This index is always guaranteed to be within the bounds of the pseudoinstruction list due to the dependent type placed on the function.
306No decoding is required.
307We then proceed by case analysis over the pseudoinstruction, reusing the code for object code for all instructions present in the MCS-51's instruction set.
308For all newly introduced pseudoinstructions, we simply translate labels to concrete addresses before behaving as a `real' instruction.
309
310We do not perform any kind of symbolic execution, wherein data is the disjoint union of bytes and addresses, with addresses kept opaque and immutable.
311Labels are immediately translated to concrete addresses, and registers and memory locations only ever contain bytes, never labels.
312As a consequence, we allow the programmer to mangle, change and generally adjust addresses as they want, under the proviso that the translation process may not be able to preserve the semantics of programs that do this.
313The only limitation introduced by this approach is that the size of
314assembly programs is bounded by $2^16$.
315This will be further discussed in Subsection~\ref{subsect.total.correctness.for.well.behaved.assembly.programs}.
316
317% ---------------------------------------------------------------------------- %
318% SECTION                                                                      %
319% ---------------------------------------------------------------------------- %
320
321\subsection{The assembler}
322\label{subsect.the.assembler}
323
324Conceptually the assembler works in two passes.
325The first pass expands every pseudoinstruction into a list of machine code instructions using the function \texttt{expand\_pseudo\_instruction}.
326The second pass encodes as a list of bytes the expanded instruction list by mapping the function \texttt{assembly1} across the list, and then flattening.
327The program obtained as a list of bytes is ready to be loaded in code memory
328for execution.
329\begin{displaymath}
330\hspace{-0.5cm}
331\mbox{\fontsize{7}{9}\selectfont$[\mathtt{P_1}, \ldots \mathtt{P_n}]$} \underset{\mbox{\fontsize{7}{9}\selectfont$\mathtt{assembly}$}}{\xrightarrow{\left(P_i \underset{\mbox{\fontsize{7}{9}\selectfont$\mathtt{assembly\_1\_pseudo\_instruction}$}}{\xrightarrow{\mathtt{P_i} \xrightarrow{\mbox{\fontsize{7}{9}\selectfont$\mathtt{expand\_pseudo\_instruction}$}} \mathtt{[I^1_i, \ldots I^q_i]} \xrightarrow{\mbox{\fontsize{7}{9}\selectfont$\mathtt{~~~~~~~~assembly1^{*}~~~~~~~~}$}} \mathtt{[0110]}}} \mathtt{[0110]}\right)^{*}}} \mbox{\fontsize{7}{9}\selectfont$\mathtt{[\ldots0110\ldots]}$}
332\end{displaymath}
333The most complex of the two passes is the first, which expands pseudoinstructions and must perform the task of branch displacement~\cite{hyde:branch:2006}.
334The function \texttt{assembly\_1\_pseudo\_instruction} used in the body of the paper is essentially the composition of the two passes.
335
336The branch displacement problem refers to the task of expanding pseudojumps into their concrete counterparts, preferably as efficiently as possible.
337For instance, the MCS-51 features three unconditional jump instructions: \texttt{LJMP} and \texttt{SJMP}---`long jump' and `short jump' respectively---and an 11-bit oddity of the MCS-51, \texttt{AJMP}.
338Each of these three instructions expects arguments in different sizes and behaves in markedly different ways: \texttt{SJMP} may only perform a `local jump'; \texttt{LJMP} may jump to any address in the MCS-51's memory space and \texttt{AJMP} may jump to any address in the current memory page.
339Consequently, the size of each opcode is different, and to squeeze as much code as possible into the MCS-51's limited code memory, the smallest possible opcode that will suffice should be selected.
340
341Similarly, a conditional pseudojump must be translated potentially into a configuration of machine code instructions, depending on the distance to the jump's target.
342For example, to translate a jump to a label, a single conditional jump pseudoinstruction may be translated into a block of three real instructions as follows (here, \texttt{JZ} is `jump if accumulator is zero'):
343{\small{
344\begin{displaymath}
345\begin{array}{r@{\quad}l@{\;\;}l@{\qquad}c@{\qquad}l@{\;\;}l}
346       & \mathtt{JZ}  & \mathtt{label}                      &                 & \mathtt{JZ}   & \text{size of \texttt{SJMP} instruction} \\
347       & \ldots       &                            & \text{translates to}   & \mathtt{SJMP} & \text{size of \texttt{LJMP} instruction} \\
348\mathtt{label:} & \mathtt{MOV} & \mathtt{A}\;\;\mathtt{B}   & \Longrightarrow & \mathtt{LJMP} & \text{address of \textit{label}} \\
349       &              &                            &                 & \ldots        & \\
350       &              &                            &                 & \mathtt{MOV}  & \mathtt{A}\;\;\mathtt{B}
351\end{array}
352\end{displaymath}}}
353Here, if \texttt{JZ} fails, we fall through to the \texttt{SJMP} which jumps over the \texttt{LJMP}.
354Naturally, if \texttt{label} is `close enough', a conditional jump pseudoinstruction is mapped directly to a conditional jump machine instruction; the above translation only applies if \texttt{label} is not sufficiently local.
355
356In order to implement branch displacement it is impossible to really make the \texttt{expand\_pseudo\_instruction} function completely independent of the encoding function.
357This is due to branch displacement requiring the distance in bytes of the target of the jump.
358Moreover the standard solutions for solving the branch displacement problem find their solutions iteratively, by either starting from a solution where all jumps are long, and shrinking them when possible, or starting from a state where all jumps are short and increasing their length as needed.
359Proving the correctness of such algorithms is already quite involved and the correctness of the assembler as a whole does not depend on the `quality' of the solution found to a branch displacement problem.
360For this reason, we try to isolate the computation of a branch displacement problem from the proof of correctness for the assembler by parameterising our \texttt{expand\_pseudo\_instruction} by a `policy'.
361
362\begin{lstlisting}
363definition expand_pseudo_instruction:
364 $\forall$lookup_labels: Identifier $\rightarrow$ Word.
365 $\forall$policy.
366 $\forall$ppc: Word.
367 $\forall$lookup_datalabels: Identifier $\rightarrow$ Word.
368 $\forall$pi: pseudo_instruction.
369  list instruction := ...
370\end{lstlisting}
371Here, the functions \texttt{lookup\_labels} and \texttt{lookup\_datalabels} are the functions that map labels and datalabels to program counters respectively, both of them used in the semantics of assembly.
372The input \texttt{pi} is the pseudoinstruction to be expanded and is found at address \texttt{ppc} in the assembly program.
373The function takes \texttt{policy} as an input.
374In reality, this is a pair of functions, but for the purposes of this paper we simplify.
375The \texttt{policy} maps pseudo-program counters to program counters: the encoding of the expansion of the pseudoinstruction found at address \texttt{a} in the assembly code should be placed into code memory at address \texttt{policy(a)}.
376Of course this is possible only if the policy is correct, which means that the encoding of consecutive assembly instructions must be consecutive in code memory.
377\begin{displaymath}
378\texttt{policy}(\texttt{ppc} + 1) = \texttt{pc} + \texttt{current\_instruction\_size}
379\end{displaymath}
380Here, \texttt{current\_instruction\_size} is the size in bytes of the encoding of the expanded pseudoinstruction found at \texttt{ppc}.
381Note that the entanglement we hinted at is only partially solved in this way: the assembler code can ignore the implementation details of the algorithm that finds a policy;
382however, the algorithm that finds a policy must know the exact behaviour of the assembly program because it needs to predict the way the assembly will expand and encode pseudoinstructions, once fed with a policy.
383A companion submission to this one~\cite{boender:correctness:2012} certifies an algorithm that finds branch displacement policies for the assembler described in this paper.
384
385The \texttt{expand\_pseudo\_instruction} function uses the \texttt{policy} map to determine the size of jump required when expanding pseudojumps, computing the jump size by examining the size of the differences between program counters.
386For instance, if at address \texttt{ppc} in the assembly program we found \texttt{Jmp l} such that \texttt{lookup\_labels l = a}, if the offset \texttt{d = policy(a) - policy(ppc + 1)} is such that \texttt{d} $< 128$ then \texttt{Jmp l} is normally translated to the best local solution, the short jump \texttt{SJMP d}.
387A global best solution to the branch displacement problem, however, is not always made of locally best solutions.
388Therefore, in some circumstances, it is necessary to force the assembler to expand jumps into larger ones.
389This is achieved by another boolean-valued function such that if the function applied to \texttt{ppc} returns true then a \texttt{Jmp l} at address \texttt{ppc} is always translated to a long jump.
390An essentially identical mechanism exists for call instructions.
391
392% ---------------------------------------------------------------------------- %
393% SECTION                                                                      %
394% ---------------------------------------------------------------------------- %
395\subsection{Correctness of the assembler with respect to fetching}
396\label{subsect.total.correctness.of.the.assembler}
397Using our policies, we now work toward proving the correctness of the assembler.
398Correctness means that the assembly process never fails when provided with a correct policy and that the process does not change the semantics of a certain class of well-behaved assembly programs.
399
400The aim of this section is to prove the following informal statement: when we fetch an assembly pseudoinstruction \texttt{I} at address \texttt{ppc}, then we can fetch the expanded pseudoinstruction(s) \texttt{[J1, \ldots, Jn] = fetch\_pseudo\_instruction \ldots\ I\ ppc} from \texttt{policy ppc} in the code memory obtained by loading the assembled object code.
401This constitutes the first major step in the proof of correctness of the assembler, the next one being the simulation of \texttt{I} by \texttt{[J1, \ldots, Jn]} (see Subsection~\ref{subsect.total.correctness.for.well.behaved.assembly.programs}).
402
403The \texttt{assembly} function is given a Russell type (slightly simplified here):
404\begin{lstlisting}
405definition assembly:
406  $\forall$program: pseudo_assembly_program.
407  $\forall$policy.
408    $\Sigma$assembled: list Byte $\times$ (BitVectorTrie costlabel 16).
409      policy is correct for program $\rightarrow$
410      $\mid$program$\mid$ < $2^{16}$ $\rightarrow$ $\mid$fst assembled$\mid$ < $2^{16}$ $\wedge$
411      (policy ($\mid$program$\mid$) = $\mid$fst assembled$\mid$ $\vee$
412      (policy ($\mid$program$\mid$) = 0 $\wedge$ $\mid$fst assembled$\mid$ = $2^{16}$)) $\wedge$
413      $\forall$ppc: pseudo_program_counter. ppc < $2^{16}$ $\rightarrow$
414        let pseudo_instr := fetch from program at ppc in
415        let assembled_i := assemble pseudo_instr in
416          $\mid$assembled_i$\mid$ $\leq$ $2^{16}$ $\wedge$
417            $\forall$n: nat. n < $\mid$assembled_i$\mid$ $\rightarrow$ $\exists$k: nat.
418              nth assembled_i n = nth assembled (policy ppc + k).
419\end{lstlisting}
420In plain words, the type of \texttt{assembly} states the following.
421Suppose we are given a policy that is correct for the program we are assembling.
422Then we return a list of assembled bytes, complete with a map from program counters to cost labels, such that the following properties hold for the list of bytes.
423Under the condition that the policy is `correct' for the program and the program is fully addressable by a 16-bit word, the assembled list is also fully addressable by a 16-bit word, the policy maps the last program counter that can address the program to the last instruction of the assemble pseudoinstruction or overflows, and if we fetch from the pseudo-program counter \texttt{ppc} we get a pseudoinstruction \texttt{pi} and a new pseudo-program counter \texttt{ppc}.
424Further, assembling the pseudoinstruction \texttt{pseudo\_instr} results in a list of bytes, \texttt{assembled\_i}.
425Then, indexing into this list with any natural number \texttt{n} less than the length of \texttt{assembled\_i} gives the same result as indexing into \texttt{assembled} with \texttt{policy ppc} (the program counter pointing to the start of the expansion in \texttt{assembled}) plus \texttt{k}.
426
427Essentially the lemma above states that the \texttt{assembly} function correctly expands pseudoinstructions, and that the expanded instruction reside consecutively in memory.
428This result is lifted from lists of bytes into a result on tries of bytes (i.e. code memories), using an additional lemma: \texttt{assembly\_ok}.
429
430Lemma \texttt{fetch\_assembly} establishes that the \texttt{fetch} and \texttt{assembly1} functions interact correctly.
431The \texttt{fetch} function, as its name implies, fetches the instruction indexed by the program counter in the code memory, while \texttt{assembly1} maps a single instruction to its byte encoding:
432\begin{lstlisting}
433lemma fetch_assembly:
434 $\forall$pc: Word.
435 $\forall$i: instruction.
436 $\forall$code_memory: BitVectorTrie Byte 16.
437 $\forall$assembled: list Byte.
438  assembled = assemble i $\rightarrow$
439  let len := $\mid$assembled$\mid$ in
440  let pc_plus_len := pc + len in
441   encoding_check pc pc_plus_len assembled $\rightarrow$
442   let $\langle$instr, pc', ticks$\rangle$ := fetch pc in
443    instr = i $\wedge$ ticks = (ticks_of_instruction instr) $\wedge$ pc' = pc_plus_len.
444\end{lstlisting}
445We read \texttt{fetch\_assembly} as follows.
446Given an instruction, \texttt{i}, we first assemble the instruction to obtain \texttt{assembled}, checking that the assembled instruction was stored in code memory correctly.
447Fetching from code memory, we obtain a tuple consisting of the instruction, new program counter, and the number of ticks this instruction will take to execute.
448We finally check that the fetched instruction is the same instruction that we began with, and the number of ticks this instruction will take to execute is the same as the result returned by a lookup function, \texttt{ticks\_of\_instruction}, devoted to tracking this information.
449Or, in plainer words, assembling and then immediately fetching again gets you back to where you started.
450
451Lemma \texttt{fetch\_assembly\_pseudo} is obtained by composition of \texttt{expand\_pseudo\_instruction} and \texttt{assembly\_1\_pseudoinstruction}:
452\begin{lstlisting}
453lemma fetch_assembly_pseudo:
454 $\forall$program: pseudo_assembly_program.
455 $\forall$policy.
456 $\forall$ppc.
457 $\forall$code_memory.
458 let $\langle$preamble, instr_list$\rangle$ := program in
459 let pi := $\pi_1$ (fetch_pseudo_instruction instr_list ppc) in
460 let pc := policy ppc in
461 let instrs := expand_pseudo_instructio policy ppc pi in
462 let $\langle$l, a$\rangle$ := assembly_1_pseudoinstruction policy ppc pi in
463 let pc_plus_len := pc + l in
464  encoding_check code_memory pc pc_plus_len a $\rightarrow$
465   fetch_many code_memory pc_plus_len pc instructions.
466\end{lstlisting}
467Here, \texttt{l} is the number of machine code instructions the pseudoinstruction at hand has been expanded into.
468We assemble a single pseudoinstruction with \texttt{assembly\_1\_pseudoinstruction}, which internally calls \texttt{expand\_pseudo\_instruction}.
469The function \texttt{fetch\_many} fetches multiple machine code instructions from code memory and performs some routine checks.
470
471Intuitively, Lemma \texttt{fetch\_assembly\_pseudo} can be read as follows.
472Suppose we expand the pseudoinstruction at \texttt{ppc} with the policy, obtaining the list of machine code instructions \texttt{instrs}.
473Suppose we also assemble the pseudoinstruction at \texttt{ppc} to obtain \texttt{a}, a list of bytes.
474Then, we check with \texttt{fetch\_many} that the number of machine instructions that were fetched matches the number of instruction that \texttt{expand\_pseudo\_instruction} expanded.
475
476The final lemma in this series is \texttt{fetch\_assembly\_pseudo2} that combines the Lemma \texttt{fetch\_assembly\_pseudo} with the correctness of the functions that load object code into the processor's memory:
477\begin{lstlisting}
478lemma fetch_assembly_pseudo2:
479 $\forall$program.
480 $\mid$snd program$\mid$ $\leq$ $2^{16}$ $\rightarrow$
481 $\forall$policy.
482 policy is correct for program $\rightarrow$
483 $\forall$ppc. ppc < $\mid$snd program$\mid$ $\rightarrow$
484  let $\langle$labels, costs$\rangle$ := create_label_cost_map program in
485  let $\langle$assembled, costs'$\rangle$ := $\pi_1$ (assembly program policy) in
486  let cmem := load_code_memory assembled in
487  let $\langle$pi, newppc$\rangle$ := fetch_pseudo_instruction program ppc in
488  let instructions := expand_pseudo_instruction policy ppc pi in
489    fetch_many cmem (policy newppc) (policy ppc) instructions.
490\end{lstlisting}
491
492Here we use $\pi_1$ to project the existential witness from the Russell-typed function \texttt{assembly}.
493
494We read \texttt{fetch\_assembly\_pseudo2} as follows.
495Suppose we are given an assembly program which can be addressed by a 16-bit word and a policy that is correct for this program.
496Suppose we are able to successfully assemble an assembly program using \texttt{assembly} and produce a code memory, \texttt{cmem}.
497Then, fetching a pseudoinstruction from the pseudo-code memory at address \texttt{ppc} corresponds to fetching a sequence of instructions from the real code memory using \texttt{policy} to expand pseudoinstructions.
498The fetched sequence corresponds to the expansion, according to the policy, of the pseudoinstruction.
499
500At first, the lemma appears to immediately imply the correctness of the assembler, but this property is \emph{not} strong enough to establish that the semantics of an assembly program has been preserved by the assembly process since it does not establish the correspondence between the semantics of a pseudoinstruction and that of its expansion.
501In particular, the two semantics differ on instructions that \emph{could} directly manipulate program addresses.
502
503% ---------------------------------------------------------------------------- %
504% SECTION                                                                      %
505% ---------------------------------------------------------------------------- %
506\subsection{Correctness for `well-behaved' assembly programs}
507\label{subsect.total.correctness.for.well.behaved.assembly.programs}
508
509The traditional approach to verifying the correctness of an assembler is to treat memory addresses as opaque structures that cannot be modified.
510Memory is represented as a map from opaque addresses to the disjoint union of data and opaque addresses---addresses are kept opaque to prevent their possible `semantics breaking' manipulation by assembly programs:
511\begin{displaymath}
512\mathtt{Mem} : \mathtt{Addr} \rightarrow \mathtt{Bytes} + \mathtt{Addr} \qquad \llbracket - \rrbracket : \mathtt{Instr} \rightarrow \mathtt{Mem} \rightarrow \mathtt{option\ Mem}
513\end{displaymath}
514The semantics of a pseudoinstruction, $\llbracket - \rrbracket$, is given as a possibly failing function from pseudoinstructions and memory spaces to new memory spaces.
515The semantic function proceeds by case analysis over the operands of a given instruction, failing if either operand is an opaque address, or otherwise succeeding, updating memory.
516\begin{gather*}
517\llbracket \mathtt{ADD\ @A1\ @A2} \rrbracket^\mathtt{M} = \begin{cases}
518                                                              \mathtt{Byte\ b1},\ \mathtt{Byte\ b2} & \rightarrow \mathtt{Some}(\mathtt{M}\ \text{with}\ \mathtt{b1} + \mathtt{b2}) \\
519                                                              -,\ \mathtt{Addr\ a} & \rightarrow \mathtt{None} \\
520                                                              \mathtt{Addr\ a},\ - & \rightarrow \mathtt{None}
521                                                            \end{cases}
522\end{gather*}
523In this paper we take a different approach, tracing memory locations (and accumulators) that contain memory addresses.
524We prove that only those assembly programs that use addresses in `safe' ways have their semantics preserved by the assembly process---a sort of dynamic type system sitting atop memory.
525In principle this approach allows us to introduce some permitted \emph{benign} manipulations of addresses that the traditional approach, using opaque addresses, cannot handle, therefore expanding the set of input programs that can be assembled correctly.
526This approach seems similar to one taken by Tuch \emph{et al}~\cite{tuch:types:2007} for reasoning about low-level C code.
527
528Our analogue of the semantic function above is merely a wrapper around the function that implements the semantics of machine code, paired with a function that keeps track of addresses.
529The semantics of pseudo- and machine code are then essentially shared.
530The only thing that changes at the assembly level is the presence of the new tracking function.
531
532However, with this approach we must detect (at run time) programs that manipulate addresses in well-behaved ways, according to some approximation of well-behavedness.
533We use an \texttt{internal\_pseudo\_address\_map} to trace addresses of code memory addresses in internal RAM:
534\begin{lstlisting}
535definition address_entry := upper_lower $\times$ Byte.
536
537definition internal_pseudo_address_map :=
538  (BitVectorTrie address_entry 7) $\times$ (BitVectorTrie address_entry 7)
539    $\times$ (option address_entry).
540\end{lstlisting}
541Here, \texttt{upper\_lower} is a type isomorphic to the booleans denoting whether a byte value is the upper or lower byte of some 16-bit address.
542
543The implementation of \texttt{internal\_pseudo\_address\_map} is complicated by some peculiarities of the MCS-51's instruction set.
544Note here that all addresses are 16 bit words, but are stored (and manipulated) as 8 bit bytes.
545All \texttt{MOV} instructions in the MCS-51 must use the accumulator \texttt{A} as an intermediary, moving a byte at a time.
546The third component of \texttt{internal\_pseudo\_address\_map} therefore states whether the accumulator currently holds a piece of an address, and if so, whether it is the upper or lower byte of the address (using the \texttt{upper\_lower} flag) complete with the corresponding source address in full.
547The first and second components, on the other hand, performs a similar task for the higher and lower external RAM.
548Again, we use our \texttt{upper\_lower} flag to describe whether a byte is the upper or lower component of a 16-bit address.
549
550The \texttt{low\_internal\_ram\_of\_pseudo\_low\_internal\_ram} function converts the lower internal RAM of a \texttt{PseudoStatus} into the lower internal RAM of a \texttt{Status}.
551A similar function exists for high internal RAM.
552Note that both RAM segments are indexed using addresses 7-bits long:
553\begin{lstlisting}
554definition low_internal_ram_of_pseudo_low_internal_ram:
555 internal_pseudo_address_map $\rightarrow$ policy $\rightarrow$ BitVectorTrie Byte 7
556  $\rightarrow$ BitVectorTrie Byte 7.
557\end{lstlisting}
558
559Next, we are able to translate \texttt{PseudoStatus} records into \texttt{Status} records using \texttt{status\_of\_pseudo\_status}.
560Translating a \texttt{PseudoStatus}'s code memory requires we expand pseudoinstructions and then assemble to obtain a trie of bytes.
561This never fails, provided that our policy is correct:
562\begin{lstlisting}
563definition status_of_pseudo_status:
564 internal_pseudo_address_map $\rightarrow$ $\forall$pap. $\forall$ps: PseudoStatus pap.
565 $\forall$policy. Status (code_memory_of_pseudo_assembly_program pap policy)
566\end{lstlisting}
567
568The \texttt{next\_internal\_pseudo\_address\_map} function is responsible for run time monitoring of the behaviour of assembly programs, in order to detect well-behaved ones.
569It returns a map that traces memory addresses in internal RAM after execution of the next pseudoinstruction, failing when the instruction tampers with memory addresses in unanticipated (but potentially correct) ways.
570It thus decides the membership of a strict subset of the set of well-behaved programs.
571\begin{lstlisting}
572definition next_internal_pseudo_address_map: internal_pseudo_address_map $\rightarrow$
573 $\forall$cm. (Identifier $\rightarrow$ PseudoStatus cm $\rightarrow$ Word) $\rightarrow$ $\forall$s: PseudoStatus cm.
574   program_counter s < $2^{16}$ $\rightarrow$ option internal_pseudo_address_map
575\end{lstlisting}
576If we wished to allow `benign manipulations' of addresses, it would be this function that needs to be changed.
577Note we once again use dependent types to ensure that program counters are properly within bounds.
578The third argument is a function that resolves the concrete address of a label.
579
580The function \texttt{ticks\_of0} computes how long---in clock cycles---a pseudoinstruction will take to execute when expanded in accordance with a given policy.
581The function returns a pair of natural numbers, needed for recording the execution times of each branch of a conditional jump.
582\begin{lstlisting}
583definition ticks_of0:
584 pseudo_assembly_program $\rightarrow$ (Identifier $\rightarrow$ Word) $\rightarrow$ $\forall$policy. Word $\rightarrow$
585   pseudo_instruction $\rightarrow$ nat $\times$ nat
586\end{lstlisting}
587An additional function, \texttt{ticks\_of}, is merely a wrapper around this function.
588
589Finally, we are able to state and prove our main theorem, relating the execution of a single assembly instruction and the execution of (possibly) many machine code instructions, as long as we are able to track memory addresses properly:
590\begin{lstlisting}
591theorem main_thm:
592 $\forall$M, M': internal_pseudo_address_map.
593 $\forall$program: pseudo_assembly_program.
594 $\forall$program_in_bounds: $\mid$program$\mid$ $\leq$ $2^{16}$.
595 let maps := create_label_cost_map program in
596 let addr_of := ... in
597 program is well labelled $\rightarrow$
598 $\forall$policy. policy is correct for program.
599 $\forall$ps: PseudoStatus program. ps < $\mid$program$\mid$.
600  next_internal_pseudo_address_map M program ... = Some M' $\rightarrow$
601   $\exists$n. execute n (status_of_pseudo_status M ps policy) =
602    status_of_pseudo_status M'
603      (execute_1_pseudo_instruction program
604       (ticks_of program ($\lambda$id. addr_of id ps) policy) ps) policy.
605\end{lstlisting}
606The statement is standard for forward simulation, but restricted to \texttt{PseudoStatuses} \texttt{ps} whose next instruction to be executed is well-behaved with respect to the \texttt{internal\_pseudo\_address\_map} \texttt{M}.
607We explicitly require proof that the policy is correct, the program is well-labelled (i.e. no repeated labels, etc.) and the pseudo-program counter is in the program's bounds.
608Theorem \texttt{main\_thm} establishes the correctness of the assembly process and can be lifted to the forward simulation of an arbitrary number of well-behaved steps on the assembly program.
609
610% ---------------------------------------------------------------------------- %
611% SECTION                                                                      %
612% ---------------------------------------------------------------------------- %
613\section{Conclusions}
614\label{sect.conclusions}
615
616We are proving the correctness of an assembler for MCS-51 assembly language.
617Our assembly language features labels, arbitrary conditional and unconditional jumps to labels, global data and instructions for moving this data into the MCS-51's single 16-bit register.
618Expanding these pseudoinstructions into machine code instructions is not trivial, and the proof that the assembly process is `correct', in that the semantics of a subset of assembly programs are not changed is complex.
619
620The formalisation is a component of CerCo which aims to produce a verified concrete complexity preserving compiler for a large subset of the C language.
621The verified assembler, complete with the underlying formalisation of the semantics of MCS-51 machine code, will form the bedrock layer upon which the rest of CerCo will build its verified compiler platform.
622
623It is interesting to compare our work to an `industrial grade' assembler for the MCS-51: SDCC~\cite{sdcc:2011}.
624SDCC is the only open source C compiler that targets the MCS-51 instruction set.
625It appears that all pseudojumps in SDCC assembly are expanded to \texttt{LJMP} instructions, the worst possible jump expansion policy from an efficiency point of view.
626Note that this policy is the only possible policy \emph{in theory} that can preserve the semantics of an assembly program during the assembly process.
627However, this comes at the expense of assembler completeness: the generated program may be too large to fit into code memory.
628In this respect, there is a trade-off between the completeness of the assembler and the efficiency of the assembled program.
629The definition and proof of a terminating, correct jump expansion policy is described in a companion publication to this one~\cite{boender:correctness:2012}.
630
631Aside from their application in verified compiler projects such as CerCo, CompCert~\cite{leroy:formally:2009} and CompCertTSO~\cite{sevcik:relaxed-memory:2011}, verified assemblers such as ours could also be applied to the verification of operating system kernels.
632Of particular note is the verified seL4 kernel~\cite{klein:sel4:2009}.
633This verification explicitly assumes the existence of, amongst other things, a trustworthy assembler and compiler.
634
635CompCert, CompCertTSO and the seL4 formalisation assume the existence of `trustworthy' assemblers.
636For instance, the CompCert C compiler's default backend stops at the PowerPC assembly language.
637The observation that an optimising assembler cannot preserve the semantics of every assembly program may have consequences for these projects (though in the case of CompCertTSO, targetting a multiprocessor, what exactly constitutes the subset of `good programs' may not be entirely clear).
638If CompCert chooses to assume the existence of an optimising assembler, then care should be made to ensure that any assembly program produced by the CompCert compiler falls into the subset of programs that can have their semantics preserved by an optimising assembler.
639
640Our formalisation exploits dependent types in different ways and for multiple purposes.
641The first purpose is to reduce potential errors in the formalisation of the microprocessor.
642Dependent types are used to constrain the size of bitvectors and tries that represent memory quantities and memory areas respectively.
643They are also used to simulate polymorphic variants in Matita, in order to provide precise typings to various functions expecting only a subset of all possible addressing modes that the MCS-51 offers.
644Polymorphic variants nicely capture the absolutely unorthogonal instruction set of the MCS-51 where every opcode must accept its own subset of the 11 addressing mode of the processor.
645
646The second purpose is to single out sources of incompleteness.
647By abstracting our functions over the dependent type of correct policies, we were able to manifest the fact that the compiler never refuses to compile a program where a correct policy exists.
648This also allowed to simplify the initial proof by dropping lemmas establishing that one function fails if and only if some previous function does so.
649
650Finally, dependent types, together with Matita's liberal system of coercions, allow us to simulate almost entirely in user space the proof methodology `Russell' of Sozeau~\cite{sozeau:subset:2006}.
651Not every proof has been carried out in this way: we only used this style to prove that a function satisfies a specification that only involves that function in a significant way.
652It would not be natural to see the proof that fetch and assembly commute as the specification of one of the two functions.
653%\paragraph{Related work}
654
655% piton
656We are not the first to consider the correctness of an assembler for a non-trivial assembly language.
657The most impressive piece of work in this domain is Piton~\cite{moore:piton:1996}, a stack of verified components, written and verified in ACL2, ranging from a proprietary FM9001 microprocessor verified at the gate level, to assemblers and compilers for two high-level languages---Lisp and $\mu$Gypsy~\cite{moore:grand:2005}.
658
659% jinja
660Klein and Nipkow consider a Java-like programming language, Jinja~\cite{klein:machine:2006}.
661They provide a compiler, virtual machine and operational semantics for the programming language and virtual machine, and prove that their compiler is semantics and type preserving.
662
663Though other verified assemblers exist in the literature what sets our work apart from that above is our attempt to optimise the generated machine code.
664This complicates any formalisation effort as an attempt at the best possible selection of machine instructions must be made---especially important on devices with limited code memory.
665Care must be taken to ensure that the time properties of an assembly program are not modified by assembly lest we affect the semantics of any program employing the MCS-51's I/O facilities.
666This is only possible by inducing a cost model on the source code from the optimisation strategy and input program.
667%\paragraph{Resources}
668
669Our source files are available at~\url{http://cerco.cs.unibo.it}.
670We assumed several properties of `library functions', e.g. modular arithmetic and datastructure manipulation.
671We axiomatised various small functions needed to complete the main theorems, as well as some `routine' proof obligations of the theorems themselves, in focussing on the main meat of the theorems.
672We believe that the proof strategy is sound and that all axioms can be closed, up to minor bugs that should have local fixes that do not affect the global proof strategy.
673
674The development, including the definition of the executable semantics of the MCS-51, is spread across 29 files, with around 18,500 lines of Matita source.
675The bulk of the proof is contained in a series of files, \texttt{AssemblyProof.ma}, \texttt{AssemblyProofSplit.ma} and \texttt{AssemblyProofSplitSplit.ma} consisting of approximately 4500 lines of Matita source.
676Numerous other lines of proofs are spread all over the development because of dependent types and the Russell proof style, which does not allow one to separate the code from the proofs.
677The low ratio between source lines and the number of lines of proof is unusual, but justified by the fact that the pseudo-assembly and the assembly language share most constructs and large swathes of the semantics are shared.
678Many lines of code are required to describe the complex semantics of the processor, but for the shared cases the proof of preservation of the semantics is essentially trivial.
679
680\bibliography{cpp-2012-asm.bib}
681
682\end{document}\renewcommand{\verb}{\lstinline}
683\def\lstlanguagefiles{lst-grafite.tex}
684\lstset{language=Grafite}
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