# Menhir Reference Manual (version 20200624) ### François Pottier and Yann Régis-Gianas INRIA {Francois.Pottier, Yann.Regis-Gianas}@inria.fr

## Contents

## 1 Foreword

Menhir is a parser generator. It turns high-level grammar specifications, decorated with semantic actions expressed in the OCaml programming language [18], into parsers, again expressed in OCaml. It is based on Knuth’s LR(1) parser construction technique [15]. It is strongly inspired by its precursors: yacc [11], ML-Yacc [22], and ocamlyacc [18], but offers a large number of minor and major improvements that make it a more modern tool.

This brief reference manual explains how to use Menhir. It does not attempt to explain context-free grammars, parsing, or the LR technique. Readers who have never used a parser generator are encouraged to read about these ideas first [1,2,8]. They are also invited to have a look at the demos directory in Menhir’s distribution.

Potential users of Menhir should be warned that Menhir’s feature set is not completely stable. There is a tension between preserving a measure of compatibility with ocamlyacc, on the one hand, and introducing new ideas, on the other hand. Some aspects of the tool, such as the error handling mechanism, are still potentially subject to incompatible changes: for instance, in the future, the current error handling mechanism (which is based on the error token, see §10) could be removed and replaced with an entirely different mechanism.

There is room for improvement in the tool and in this reference manual. Bug reports and suggestions are welcome!

## 2 Usage

Menhir is invoked as follows:

> menhir option …option filename …filename

Each of the file names must end with .mly (unless --coq is used, in which case it must end with .vy) and denotes a partial grammar specification. These partial grammar specifications are joined (§5.1) to form a single, self-contained grammar specification, which is then processed. The following optional command line switches allow controlling many aspects of the process.

--base basename. This switch controls the base name of the .ml and .mli files that are produced. That is, the tool will produce files named basename.ml and basename.mli. Note that basename can contain occurrences of the / character, so it really specifies a path and a base name. When only one filename is provided on the command line, the default basename is obtained by depriving filename of its final .mly suffix. When multiple file names are provided on the command line, no default base name exists, so that the --base switch must be used.

--cmly. This switch causes Menhir to produce a .cmly file in addition to its normal operation. This file contains a (binary-form) representation of the grammar and automaton (see §13.1).

--comment. This switch causes a few comments to be inserted into the OCaml code that is written to the .ml file.

--compare-errors filename1 --compare-errors filename2. Two such switches must always be used in conjunction so as to specify the names of two .messages files, filename1 and filename2. Each file is read and internally translated to a mapping of states to messages. Menhir then checks that the left-hand mapping is a subset of the right-hand mapping. This feature is typically used in conjunction with --list-errors to check that filename2 is complete (that is, covers all states where an error can occur). For more information, see §11.

--compile-errors filename. This switch causes Menhir to read the file filename, which must obey the .messages file format, and to compile it to an OCaml function that maps a state number to a message. The OCaml code is sent to the standard output channel. At the same time, Menhir checks that the collection of input sentences in the file filename is correct and irredundant. For more information, see §11.

--coq. This switch causes Menhir to produce Coq code. See §12.

--coq-lib-path path. This switch allows specifying under what name (or path) the Coq support library MenhirLib is known to Coq. When Menhir runs in --coq mode, the generated parser contains references to several modules in this library. This path is used to qualify these references. Its default value is MenhirLib.

--coq-lib-no-path. This switch indicates that references to the Coq library MenhirLib should not be qualified. This was the default behavior of Menhir prior to 2018/05/30. This switch is provided for compatibility, but normally should not be used.

--coq-no-actions. (Used in conjunction with --coq.) This switch causes the semantic actions present in the .vy file to be ignored and replaced with `tt`, the unique inhabitant of Coq’s `unit` type. This feature can be used to test the Coq back-end with a standard grammar, that is, a grammar that contains OCaml semantic actions. Just rename the file from .mly to .vy and set this switch.

--coq-no-complete. (Used in conjunction with --coq.) This switch disables the generation of the proof of completeness of the parser (§12). This can be necessary because the proof of completeness is possible only if the grammar has no conflict (not even a benign one, in the sense of §6.1). This can be desirable also because, for a complex grammar, completeness may require a heavy certificate and its validation by Coq may take time.

--depend. See §14.

--dump. This switch causes a description of the automaton to be written to the file basename.automaton.

--echo-errors filename. This switch causes Menhir to read the .messages file filename and to produce on the standard output channel just the input sentences. (That is, all messages, blank lines, and comments are filtered out.) For more information, see §11.

--explain. This switch causes conflict explanations to be written to the file basename.conflicts. See also §6.

--external-tokens T. This switch causes the definition of the token type to be omitted in basename.ml and basename.mli. Instead, the generated parser relies on the type T.token, where T is an OCaml module name. It is up to the user to define module T and to make sure that it exports a suitable token type. Module T can be hand-written. It can also be automatically generated out of a grammar specification using the --only-tokens switch.

--fixed-exception. This switch causes the exception Error to be internally defined as a synonym for Parsing.Parse_error. This means that an exception handler that catches Parsing.Parse_error will also catch the generated parser’s Error. This helps increase Menhir’s compatibility with ocamlyacc. There is otherwise no reason to use this switch.

--graph. This switch causes a description of the grammar’s dependency graph to be written to the file basename.dot. The graph’s vertices are the grammar’s nonterminal symbols. There is a directed edge from vertex A to vertex B if the definition of A refers to B. The file is in a format that is suitable for processing by the graphviz toolkit.

--infer, --infer-write-query, --infer-read-reply. See §14.

--inspection. This switch requires --table. It causes Menhir to generate not only the monolithic and incremental APIs (§9.1, §9.2), but also the inspection API (§9.3). Activating this switch causes a few more tables to be produced, resulting in somewhat larger code size.

--interpret. This switch causes Menhir to act as an interpreter, rather than as a compiler. No OCaml code is generated. Instead, Menhir reads sentences off the standard input channel, parses them, and displays outcomes. This switch can be usefully combined with --trace. For more information, see §8.

--interpret-error. This switch is analogous to --interpret, except Menhir expects every sentence to cause an error on its last token, and displays information about the state in which the error is detected, in the .messages file format. For more information, see §11.

--interpret-show-cst. This switch, used in conjunction with --interpret, causes Menhir to display a concrete syntax tree when a sentence is successfully parsed. For more information, see §8.

--list-errors. This switch causes Menhir to produce (on the standard output channel) a complete list of input sentences that cause an error, in the .messages file format. For more information, see §11.

--log-automaton level. When level is nonzero, this switch causes some information about the automaton to be logged to the standard error channel.

--log-code level. When level is nonzero, this switch causes some information about the generated OCaml code to be logged to the standard error channel.

--log-grammar level. When level is nonzero, this switch causes some information about the grammar to be logged to the standard error channel. When level is 2, the nullable, FIRST, and FOLLOW tables are displayed.

--no-dollars. This switch disallows the use of positional keywords of the form $i.

--no-inline. This switch causes all %inline keywords in the grammar specification to be ignored. This is especially useful in order to understand whether these keywords help solve any conflicts.

--no-stdlib. This switch instructs Menhir to not use its standard library (§5.4).

--ocamlc command. See §14.

--ocamldep command. See §14.

--only-preprocess. This switch causes the grammar specifications to be transformed up to the point where the automaton’s construction can begin. The grammar specifications whose names are provided on the command line are joined (§5.1); all parameterized nonterminal symbols are expanded away (§5.2); type inference is performed, if --infer is enabled; all nonterminal symbols marked %inline are expanded away (§5.3). This yields a single, monolithic grammar specification, which is printed on the standard output channel.

--only-tokens. This switch causes the %token declarations in the grammar specification to be translated into a definition of the token type, which is written to the files basename.ml and basename.mli. No code is generated. This is useful when a single set of tokens is to be shared between several parsers. The directory demos/calc-two contains a demo that illustrates the use of this switch.

--raw-depend. See §14.

--stdlib directory. This switch exists only for backwards compatibility and is ignored. It may be removed in the future.

--strict. This switch causes several warnings about the grammar and about the automaton to be considered errors. This includes warnings about useless precedence declarations, non-terminal symbols that produce the empty language, unreachable non-terminal symbols, productions that are never reduced, conflicts that are not resolved by precedence declarations, and end-of-stream conflicts.

--suggest-*. See §14.

--table. This switch causes Menhir to use its table-based back-end, as opposed to its (default) code-based back-end. When --table is used, Menhir produces significantly more compact and somewhat slower parsers. See §16 for a speed comparison.

The table-based back-end produces rather compact tables, which are analogous to those produced by yacc, bison, or ocamlyacc. These tables are not quite stand-alone: they are exploited by an interpreter, which is shipped as part of the support library MenhirLib. For this reason, when --table is used, MenhirLib must be made visible to the OCaml compilers, and must be linked into your executable program. The --suggest-* switches, described above, help do this.

The code-based back-end compiles the LR automaton directly into a nest of mutually recursive OCaml functions. In that case, MenhirLib is not required.

The incremental API (§9.2) and the inspection API (§9.3) are made available only by the table-based back-end.

--timings. This switch causes internal timing information to be sent to the standard error channel.

--timings-to filename. This switch causes internal timing information to be written to the file filename.

--trace. This switch causes tracing code to be inserted into the generated parser, so that, when the parser is run, its actions are logged to the standard error channel. This is analogous to ocamlrun’s p=1 parameter, except this switch must be enabled at compile time: one cannot selectively enable or disable tracing at runtime.

--unused-precedence-levels. This switch suppresses all warnings about useless %left, %right, %nonassoc and %prec declarations.

--unused-token symbol. This switch suppresses the warning that is normally emitted when Menhir finds that the terminal symbol symbol is unused.

--unused-tokens. This switch suppresses all of the warnings that are normally emitted when Menhir finds that some terminal symbols are unused.

--update-errors filename. This switch causes Menhir to read the .messages file filename and to produce on the standard output channel a new .messages file that is identical, except the auto-generated comments have been re-generated. For more information, see §11.

--version. This switch causes Menhir to print its own version number and exit.

## 3 Lexical conventions

A semicolon character (;) may appear after a declaration (§4.1).

An old-style rule (§4.2) may be terminated with a semicolon. Also, within an old-style rule, each producer (§4.2.3) may be terminated with a semicolon.

A new-style rule (§4.3) must not be terminated with a semicolon. Within such a rule, the elements of a sequence must be separated with semicolons.

Semicolons are not allowed to appear anywhere except in the places mentioned above. This is in contrast with ocamlyacc, which views semicolons as insignificant, just like whitespace.

Identifiers (id) coincide with OCaml identifiers, except they are not allowed to contain the quote (’) character. Following OCaml, identifiers that begin with a lowercase letter (lid) or with an uppercase letter (uid) are distinguished.

A quoted identifier qid is a string enclosed in double quotes. Such a string cannot contain a double quote or a backslash. Quoted identifiers are used as token aliases (§4.1.3).

Comments are C-style (surrounded with /* and */, cannot be nested), C++-style (announced by // and extending until the end of the line), or OCaml-style (surrounded with (* and *), can be nested). Of course, inside OCaml code, only OCaml-style comments are allowed.

OCaml type expressions are surrounded with < and >. Within such expressions, all references to type constructors (other than the built-in list, option, etc.) must be fully qualified.

## 4 Syntax of grammar specifications

> specification | ::= | declaration … declaration %% rule … rule [ %% OCaml code ]
> declaration | ::= | %{ OCaml code %}
> |  | %parameter < uid : OCaml module type >
> |  | %token [ < OCaml type > ] uid [ qid ] … uid [ qid ]
> |  | %nonassoc uid … uid
> |  | %left uid … uid
> |  | %right uid … uid
> |  | %type < OCaml type > lid … lid
> |  | %start [ < OCaml type > ] lid … lid
> |  | %attribute actual … actual attribute … attribute
> |  | % attribute
> |  | %on_error_reduce lid … lid
> attribute | ::= | [@ name payload ]
> old syntax — rule | ::= | [ %public ] [ %inline ] lid [ ( id, …, id ) ] : [ | ] group | … | group
> group | ::= | production | … | production { OCaml code } [ %prec id ]
> production | ::= | producer … producer [ %prec id ]
> producer | ::= | [ lid = ] actual
> actual | ::= | id [ ( actual, …, actual ) ]
> |  | actual ( ? ∣ + ∣ * )
> |  | group | … | group
> new syntax — rule | ::= | [ %public ] let lid [ ( id, …, id ) ] ( := ∣ == ) expression
> expression | ::= | [ | ] expression | … | expression
> |  | [ pattern = ] expression ; expression
> |  | id [ ( expression , …, expression ) ]
> |  | expression ( ? ∣ + ∣ * )
> |  | { OCaml code } [ %prec id ]
> |  | < OCaml id > [ %prec id ]
> pattern | ::= | lid ∣ _ ∣ ~ ∣ ( pattern , …, pattern )

> Figure 1: Syntax of grammar specifications

The syntax of grammar specifications appears in Figure 1. The places where attributes can be attached are not shown; they are documented separately (§13.2). A grammar specification begins with a sequence of declarations (§4.1), ended by a mandatory %% keyword. Following this keyword, a sequence of rules is expected. Each rule defines a nonterminal symbol lid, whose name must begin with a lowercase letter. A rule is expressed either in the “old syntax” (§4.2) or in the “new syntax” (§4.3), which is slightly more elegant and powerful.

### 4.1 Declarations

#### 4.1.1 Headers

A header is a piece of OCaml code, surrounded with %{ and %}. It is copied verbatim at the beginning of the .ml file. It typically contains OCaml open directives and function definitions for use by the semantic actions. If a single grammar specification file contains multiple headers, their order is preserved. However, when two headers originate in distinct grammar specification files, the order in which they are copied to the .ml file is unspecified.

It is important to note that the header is copied by Menhir only to the .ml file, not to the .mli file. Therefore, it should not contain declarations that affect the meaning of the types that appear in the .mli file. Here are two problems that people commonly run into:

- Placing an open directive that is required for a %type declaration to make sense. For instance, writing `open Foo` in the header and declaring `%type<t> bar`, where the type `t` is defined in the module `Foo`, will not work. You must write `%type<Foo.t> bar`.
- Declaring a module alias that affects a (declared or inferred) type. For instance, writing `module F = Foo` in the header and declaring `%type<Foo.t> bar` may not work (from 2020/05/25 on). The reason is, OCaml may infer that the symbol `bar` has type `F.t`, and Menhir will rely on this information without realizing that `F` is a local name, so in the end, the .mli file contains a reference to `F.t` that does not make sense.

#### 4.1.2 Parameters

A declaration of the form:

> %parameter < uid : OCaml module type >

causes the entire parser to become parameterized over the OCaml module uid, that is, to become an OCaml functor. The directory demos/calc-param contains a demo that illustrates the use of this switch.

If a single specification file contains multiple %parameter declarations, their order is preserved, so that the module name uid introduced by one declaration is effectively in scope in the declarations that follow. When two %parameter declarations originate in distinct grammar specification files, the order in which they are processed is unspecified. Last, %parameter declarations take effect before %{ … %}, %token, %type, or %start declarations are considered, so that the module name uid introduced by a %parameter declaration is effectively in scope in all %{ … %}, %token, %type, or %start declarations, regardless of whether they precede or follow the %parameter declaration. This means, in particular, that the side effects of an OCaml header are observed only when the functor is applied, not when it is defined.

#### 4.1.3 Tokens

A declaration of the form:

> %token [ < OCaml type > ] uid1 [ qid1 ] … uidn [ qidn ]

defines the identifiers uid1, …, uidn as tokens, that is, as terminal symbols in the grammar specification and as data constructors in the token type.

If an OCaml type t is present, then these tokens are considered to carry a semantic value of type t, otherwise they are considered to carry no semantic value.

If a quoted identifier qidi is present, then it is considered an alias for the terminal symbol uidi. (This feature, known as “token aliases”, is borrowed from Bison.) Throughout the grammar, the quoted identifier qidi is then synonymous with the identifier uidi. For example, if one declares:

    %token PLUS "+"

then the quoted identifier "+" stands for the terminal symbol PLUS throughout the grammar. An example of the use of token aliases appears in the directory demos/calc-alias. Token aliases can be used to improve the readability of a grammar. One must keep in mind, however, that they are just syntactic sugar: they are not interpreted in any way by Menhir or conveyed to tools like ocamllex. They could be considered confusing by a reader who mistakenly believes that they are interpreted as string literals.

#### 4.1.4 Priority and associativity

A declaration of one of the following forms:

> %nonassoc uid1 … uidn
> %left uid1 … uidn
> %right uid1 … uidn

assigns both a priority level and an associativity status to the symbols uid1, …, uidn. The priority level assigned to uid1, …, uidn is not defined explicitly: instead, it is defined to be higher than the priority level assigned by the previous %nonassoc, %left, or %right declaration, and lower than that assigned by the next %nonassoc, %left, or %right declaration. The symbols uid1, …, uidn can be tokens (defined elsewhere by a %token declaration) or dummies (not defined anywhere). Both can be referred to as part of %prec annotations. Associativity status and priority levels allow shift/reduce conflicts to be silently resolved (§6).

#### 4.1.5 Types

A declaration of the form:

> %type < OCaml type > lid1 … lidn

assigns an OCaml type to each of the nonterminal symbols lid1, …, lidn. For start symbols, providing an OCaml type is mandatory, but is usually done as part of the %start declaration. For other symbols, it is optional. Providing type information can improve the quality of OCaml’s type error messages.

A %type declaration may concern not only a nonterminal symbol, such as, say, expression, but also a fully applied parameterized nonterminal symbol, such as list(expression) or separated_list(COMMA, option(expression)).

The types provided as part of %type declarations are copied verbatim to the .ml and .mli files. In contrast, headers (§4.1.1) are copied to the .ml file only. For this reason, the types provided as part of %type declarations must make sense both in the presence and in the absence of these headers. They should typically be fully qualified types.

#### 4.1.6 Start symbols

A declaration of the form:

> %start [ < OCaml type > ] lid1 … lidn

declares the nonterminal symbols lid1, …, lidn to be start symbols. Each such symbol must be assigned an OCaml type either as part of the %start declaration or via separate %type declarations. Each of lid1, …, lidn becomes the name of a function whose signature is published in the .mli file and that can be used to invoke the parser.

#### 4.1.7 Attribute declarations

Attribute declarations of the form %attribute actual … actual attribute … attribute and % attribute are explained in §13.2.

#### 4.1.8 Extra reductions on error

A declaration of the form:

> %on_error_reduce lid1 … lidn

marks the nonterminal symbols lid1, …, lidn as potentially eligible for reduction when an invalid token is found. This may cause one or more extra reduction steps to be performed before the error is detected.

More precisely, this declaration affects the automaton as follows. Let us say that a production lid → … is “reducible on error” if its left-hand symbol lid appears in a %on_error_reduce declaration. After the automaton has been constructed and after any conflicts have been resolved, in every state s, the following algorithm is applied:

1. Construct the set of all productions that are ready to be reduced in state s and are reducible on error;
2. Test if one of them, say p, has higher “on-error-reduce-priority” than every other production in this set;
3. If so, in state s, replace every error action with a reduction of the production p. (In other words, for every terminal symbol t, if the action table says: “in state s, when the next input symbol is t, fail”, then this entry is replaced with: “in state s, when the next input symbol is t, reduce production p”.)

If step 3 above is executed in state s, then an error can never be detected in state s, since all error actions in state s are replaced with reduce actions. Error detection is deferred: at least one reduction takes place before the error is detected. It is a “spurious” reduction: in a canonical LR(1) automaton, it would not take place.

An %on_error_reduce declaration does not affect the language that is accepted by the automaton. It does not affect the location where an error is detected. It is used to control in which state an error is detected. If used wisely, it can make errors easier to report, because they are detected in a state for which it is easier to write an accurate diagnostic message (§11.3).

Like a %type declaration, an %on_error_reduce declaration may concern not only a nonterminal symbol, such as, say, expression, but also a fully applied parameterized nonterminal symbol, such as list(expression) or separated_list(COMMA, option(expression)).

The “on-error-reduce-priority” of a production is that of its left-hand symbol. The “on-error-reduce-priority” of a nonterminal symbol is determined implicitly by the order of %on_error_reduce declarations. In the declaration %on_error_reduce lid1 … lidn, the symbols lid1, …, lidn have the same “on-error-reduce-priority”. They have higher “on-error-reduce-priority” than the symbols listed in previous %on_error_reduce declarations, and lower “on-error-reduce-priority” than those listed in later %on_error_reduce declarations.

### 4.2 Rules—old syntax

In its simplest form, a rule begins with the nonterminal symbol lid, followed by a colon character (:), and continues with a sequence of production groups (§4.2.1). Each production group is preceded with a vertical bar character (|); the very first bar is optional. The meaning of the bar is choice: the nonterminal symbol id develops to either of the production groups. We defer explanations of the keyword %public (§5.1), of the keyword %inline (§5.3), and of the optional formal parameters ( id, …, id ) (§5.2).

#### 4.2.1 Production groups

In its simplest form, a production group consists of a single production (§4.2.2), followed by an OCaml semantic action (§4.2.1) and an optional %prec annotation (§4.2.1). A production specifies a sequence of terminal and nonterminal symbols that should be recognized, and optionally binds identifiers to their semantic values.

#### Semantic actions

A semantic action is a piece of OCaml code that is executed in order to assign a semantic value to the nonterminal symbol with which this production group is associated. A semantic action can refer to the (already computed) semantic values of the terminal or nonterminal symbols that appear in the production via the semantic value identifiers bound by the production.

For compatibility with ocamlyacc, semantic actions can also refer to unnamed semantic values via positional keywords of the form $1, $2, etc. This style is discouraged. (It is in fact forbidden if --no-dollars is turned on.) Furthermore, as a positional keyword of the form $i is internally rewritten as _i, the user should not use identifiers of the form _i.

#### %prec annotations

An annotation of the form %prec id indicates that the precedence level of the production group is the level assigned to the symbol id via a previous %nonassoc, %left, or %right declaration (§4.1.4). In the absence of a %prec annotation, the precedence level assigned to each production is the level assigned to the rightmost terminal symbol that appears in it. It is undefined if the rightmost terminal symbol has an undefined precedence level or if the production mentions no terminal symbols at all. The precedence level assigned to a production is used when resolving shift/reduce conflicts (§6).

#### Multiple productions in a group

If multiple productions are present in a single group, then the semantic action and precedence annotation are shared between them. This short-hand effectively allows several productions to share a semantic action and precedence annotation without requiring textual duplication. It is legal only when every production binds exactly the same set of semantic value identifiers and when no positional semantic value keywords ($1, etc.) are used.

#### 4.2.2 Productions

A production is a sequence of producers (§4.2.3), optionally followed by a %prec annotation (§4.2.1). If a precedence annotation is present, it applies to this production alone, not to other productions in the production group. It is illegal for a production and its production group to both carry %prec annotations.

#### 4.2.3 Producers

A producer is an actual (§4.2.4), optionally preceded with a binding of a semantic value identifier, of the form lid =. The actual specifies which construction should be recognized and how a semantic value should be computed for that construction. The identifier lid, if present, becomes bound to that semantic value in the semantic action that follows. Otherwise, the semantic value can be referred to via a positional keyword ($1, etc.).

#### 4.2.4 Actuals

In its simplest form, an actual is just a terminal or nonterminal symbol id. If it is a parameterized non-terminal symbol (see §5.2), then it should be applied: id( actual, …, actual ) .

An actual may be followed with a modifier (?, +, or *). This is explained further on (see §5.2 and Figure 2).

An actual may also be an “anonymous rule”. In that case, one writes just the rule’s right-hand side, which takes the form group | … | group. (This form is allowed only as an argument in an application.) This form is expanded on the fly to a definition of a fresh non-terminal symbol, which is declared %inline. For instance, providing an anonymous rule as an argument to list:

> list ( e = expression; SEMICOLON { e } )

is equivalent to writing this:

> list ( expression_SEMICOLON )

where the non-terminal symbol expression_SEMICOLON is chosen fresh and is defined as follows:

> %inline expression_SEMICOLON:
> | e = expression; SEMICOLON { e }

### 4.3 Rules—new syntax

Please be warned that the new syntax is considered experimental and is subject to change in the future.

In its simplest form, a rule takes the form let lid := expression. Its left-hand side lid is a nonterminal symbol; its right-hand side is an expression. Such a rule defines an ordinary nonterminal symbol, while the alternate form let lid == expression defines an %inline nonterminal symbol (§5.3), that is, a macro. A rule can be preceded with the keyword %public (§5.1) and can be parameterized with a tuple of formal parameters ( id, …, id ) (§5.2). The various forms of expressions, listed in Figure 1, are:

- A choice between several expressions, [ | ] expression1 | … | expressionn. The leading bar is optional.
- A sequence of two expressions, pattern = expression1 ; expression2. The semantic value produced by expression1 is decomposed according to the pattern pattern. The OCaml variables introduced by pattern may appear in a semantic action that ends the sequence expression2.
- A sequence ~ = id1 ; expression2, which is sugar for id1 = id1 ; expression2. This is a pun.
- A sequence expression1 ; expression2, which is sugar for _ = expression1 ; expression2.
- A symbol id, possibly applied to a tuple of expressions ( expression1, …, expressionn ). It is worth noting that such an expression can form the end of a sequence: id at the end of a sequence stands for x = id ; { x } for some fresh variable x. Thus, a sequence need not end with a semantic action.
- An expression followed with ?, +, or *. This is sugar for the previous form: see §5.2 and Figure 2.
- A semantic action { OCaml code } , possibly followed with a precedence annotation %prec id. This OCaml code can refer to the variables that have been bound earlier in the sequence that this semantic action ends. These include all variables named by the user as well as all variables introduced by a ~ pattern as part of a pun. The notation $i, where i is an integer, is forbidden.
- A point-free semantic action < OCaml id >, possibly followed with a precedence annotation %prec id. The OCaml identifier id must denote a function or a data constructor. It is applied to a tuple of the variables that have been bound earlier in the sequence that this semantic action ends. Thus, < id > is sugar for { id (x1, …, xn) } , where x1, …, xn are the variables bound earlier. These include all variables named by the user as well as all variables introduced by a ~ pattern.
- An identity semantic action <>. This is sugar for < identity >, where identity is OCaml’s identity function. Therefore, it is sugar for { (x1, …, xn) } , where x1, …, xn are the variables bound earlier.

The syntax of expressions, as presented in Figure 1, seems more permissive than it really is. In reality, a choice cannot be nested inside a sequence; a sequence cannot be nested in the left-hand side of a sequence; a semantic action cannot appear in the left-hand side of a sequence. (Thus, there is a stratification in three levels: choice expressions, sequence expressions, and atomic expressions, which corresponds roughly to the stratification of rules, productions, and producers in the old syntax.) Furthermore, an expression between parentheses ( expression ) is not a valid expression. To surround an expression with parentheses, one must write either midrule ( expression ) or endrule ( expression ) ; see §5.4 and Figure 3.

When a complex expression (e.g., a choice or a sequence) is placed in parentheses, as in id ( expression ), this is equivalent to using id ( s ) , where the fresh symbol s is declared as a synonym for this expression, via the declaration let s == expression. This idiom is also known as an anonymous rule (§4.2.4).

#### Examples

As an example of a rule in the new syntax, the parameterized nonterminal symbol option, which is part of Menhir’s standard library (§5.4), can be defined as follows:

> let option(x) :=
> | { None }
> | x = x ; { Some x }

Using a pun, it can also be written as follows:

> let option(x) :=
> | { None }
> | ~ = x ; { Some x }

Using a pun and a point-free semantic action, it can also be expressed as follows:

> let option(x) :=
> | { None }
> | ~ = x ; < Some >

As another example, the parameterized symbol delimited, also part of Menhir’s standard library (§5.4), can be defined in the new syntax as follows:

> let delimited(opening, x, closing) ==
> opening ; ~ = x ; closing ; <>

The use of == indicates that this is a macro, i.e., an %inline nonterminal symbol (see §5.3). The identity semantic action <> is here synonymous with { x }.

Other illustrations of the new syntax can be found in the directories demos/calc-new-syntax and demos/calc-ast.

## 5 Advanced features

### 5.1 Splitting specifications over multiple files

#### Modules

Grammar specifications can be split over multiple files. When Menhir is invoked with multiple argument file names, it considers each of these files as a partial grammar specification, and joins these partial specifications in order to obtain a single, complete specification.

This feature is intended to promote a form a modularity. It is hoped that, by splitting large grammar specifications into several “modules”, they can be made more manageable. It is also hoped that this mechanism, in conjunction with parameterization (§5.2), will promote sharing and reuse. It should be noted, however, that this is only a weak form of modularity. Indeed, partial specifications cannot be independently processed (say, checked for conflicts). It is necessary to first join them, so as to form a complete grammar specification, before any kind of grammar analysis can be done.

This mechanism is, in fact, how Menhir’s standard library (§5.4) is made available: even though its name does not appear on the command line, it is automatically joined with the user’s explicitly-provided grammar specifications, making the standard library’s definitions globally visible.

A partial grammar specification, or module, contains declarations and rules, just like a complete one: there is no visible difference. Of course, it can consist of only declarations, or only rules, if the user so chooses. (Don’t forget the mandatory %% keyword that separates declarations and rules. It must be present, even if one of the two sections is empty.)

#### Private and public nonterminal symbols

It should be noted that joining is not a purely textual process. If two modules happen to define a nonterminal symbol by the same name, then it is considered, by default, that this is an accidental name clash. In that case, each of the two nonterminal symbols is silently renamed so as to avoid the clash. In other words, by default, a nonterminal symbol defined in module A is considered private, and cannot be defined again, or referred to, in module B.

Naturally, it is sometimes desirable to define a nonterminal symbol N in module A and to refer to it in module B. This is permitted if N is public, that is, if either its definition carries the keyword %public or N is declared to be a start symbol. A public nonterminal symbol is never renamed, so it can be referred to by modules other than its defining module.

In fact, it is permitted to split the definition of a public nonterminal symbol, over multiple modules and/or within a single module. That is, a public nonterminal symbol N can have multiple definitions, within one module and/or in distinct modules. All of these definitions are joined using the choice (|) operator. For instance, in the grammar of a programming language, the definition of the nonterminal symbol expression could be split into multiple modules, where one module groups the expression forms that have to do with arithmetic, one module groups those that concern function definitions and function calls, one module groups those that concern object definitions and method calls, and so on.

#### Tokens aside

Another use of modularity consists in placing all %token declarations in one module, and the actual grammar specification in another module. The module that contains the token definitions can then be shared, making it easier to define multiple parsers that accept the same type of tokens. (On this topic, see demos/calc-two.)

### 5.2 Parameterizing rules

A rule (that is, the definition of a nonterminal symbol) can be parameterized over an arbitrary number of symbols, which are referred to as formal parameters.

#### Example

For instance, here is the definition of the parameterized nonterminal symbol option, taken from the standard library (§5.4):

> %public option(X):
> | { None }
> | x = X { Some x }

This definition states that option(X) expands to either the empty string, producing the semantic value None, or to the string X, producing the semantic value Some x, where x is the semantic value of X. In this definition, the symbol X is abstract: it stands for an arbitrary terminal or nonterminal symbol. The definition is made public, so option can be referred to within client modules.

A client who wishes to use option simply refers to it, together with an actual parameter – a symbol that is intended to replace X. For instance, here is how one might define a sequence of declarations, preceded with optional commas:

> declarations:
> | { [] }
> | ds = declarations; option(COMMA); d = declaration { d :: ds }

This definition states that declarations expands either to the empty string or to declarations followed by an optional comma followed by declaration. (Here, COMMA is presumably a terminal symbol.) When this rule is encountered, the definition of option is instantiated: that is, a copy of the definition, where COMMA replaces X, is produced. Things behave exactly as if one had written:

> optional_comma:
> | { None }
> | x = COMMA { Some x }
> declarations:
> | { [] }
> | ds = declarations; optional_comma; d = declaration { d :: ds }

Note that, even though COMMA presumably has been declared as a token with no semantic value, writing x = COMMA is legal, and binds x to the unit value. This design choice ensures that the definition of option makes sense regardless of the nature of X: that is, X can be instantiated with a terminal symbol, with or without a semantic value, or with a nonterminal symbol.

#### Parameterization in general

In general, the definition of a nonterminal symbol N can be parameterized with an arbitrary number of formal parameters. When N is referred to within a production, it must be applied to the same number of actuals. In general, an actual is:

- either a single symbol, which can be a terminal symbol, a nonterminal symbol, or a formal parameter;
- or an application of such a symbol to a number of actuals.

For instance, here is a rule whose single production consists of a single producer, which contains several, nested actuals. (This example is discussed again in §5.4.)

> plist(X):
> | xs = loption(delimited(LPAREN, separated_nonempty_list(COMMA, X), RPAREN)) { xs }

> actual? |  | is syntactic sugar for |  | option(actual)
> actual+ |  | is syntactic sugar for |  | nonempty_list(actual)
> actual* |  | is syntactic sugar for |  | list(actual)

> Figure 2: Syntactic sugar for simulating regular expressions, also known as EBNF

Applications of the parameterized nonterminal symbols option, nonempty_list, and list, which are defined in the standard library (§5.4), can be written using a familiar, regular-expression like syntax (Figure 2).

#### Higher-order parameters

A formal parameter can itself expect parameters. For instance, here is a rule that defines the syntax of procedures in an imaginary programming language:

> procedure(list):
> | PROCEDURE ID list(formal) SEMICOLON block SEMICOLON { … }

This rule states that the token ID, which represents the name of the procedure, should be followed with a list of formal parameters. (The definitions of the nonterminal symbols formal and block are not shown.) However, because list is a formal parameter, as opposed to a concrete nonterminal symbol defined elsewhere, this definition does not specify how the list is laid out: which token, if any, is used to separate, or terminate, list elements? is the list allowed to be empty? and so on. A more concrete notion of procedure is obtained by instantiating the formal parameter list: for instance, procedure(plist), where plist is the parameterized nonterminal symbol defined earlier, is a valid application.

#### Consistency

Definitions and uses of parameterized nonterminal symbols are checked for consistency before they are expanded away. In short, it is checked that, wherever a nonterminal symbol is used, it is supplied with actual arguments in appropriate number and of appropriate nature. This guarantees that expansion of parameterized definitions terminates and produces a well-formed grammar as its outcome.

### 5.3 Inlining

It is well-known that the following grammar of arithmetic expressions does not work as expected: that is, in spite of the priority declarations, it has shift/reduce conflicts.

> %token < int > INT
> %token PLUS TIMES
> %left PLUS
> %left TIMES
> %%
> expression:
> | i = INT { i }
> | e = expression; o = op; f = expression { o e f }
> op:
> | PLUS { ( + ) }
> | TIMES { ( * ) }

The trouble is, the precedence level of the production expression → expression op expression is undefined, and there is no sensible way of defining it via a %prec declaration, since the desired level really depends upon the symbol that was recognized by op: was it PLUS or TIMES?

The standard workaround is to abandon the definition of op as a separate nonterminal symbol, and to inline its definition into the definition of expression, like this:

> expression:
> | i = INT { i }
> | e = expression; PLUS; f = expression { e + f }
> | e = expression; TIMES; f = expression { e * f }

This avoids the shift/reduce conflict, but gives up some of the original specification’s structure, which, in realistic situations, can be damageable. Fortunately, Menhir offers a way of avoiding the conflict without manually transforming the grammar, by declaring that the nonterminal symbol op should be inlined:

> expression:
> | i = INT { i }
> | e = expression; o = op; f = expression { o e f }
> %inline op:
> | PLUS { ( + ) }
> | TIMES { ( * ) }

The %inline keyword causes all references to op to be replaced with its definition. In this example, the definition of op involves two productions, one that develops to PLUS and one that expands to TIMES, so every production that refers to op is effectively turned into two productions, one that refers to PLUS and one that refers to TIMES. After inlining, op disappears and expression has three productions: that is, the result of inlining is exactly the manual workaround shown above.

In some situations, inlining can also help recover a slight efficiency margin. For instance, the definition:

> %inline plist(X):
> | xs = loption(delimited(LPAREN, separated_nonempty_list(COMMA, X), RPAREN)) { xs }

effectively makes plist(X) an alias for the right-hand side loption(…). Without the %inline keyword, the language recognized by the grammar would be the same, but the LR automaton would probably have one more state and would perform one more reduction at run time.

The %inline keyword does not affect the computation of positions (§7). The same positions are computed, regardless of where %inline keywords are placed.

If the semantic actions have side effects, the %inline keyword can affect the order in which these side effects take place. In the example of op and expression above, if for some reason the semantic action associated with op has a side effect (such as updating a global variable, or printing a message), then, by inlining op, we delay this side effect, which takes place after the second operand has been recognized, whereas in the absence of inlining it takes place as soon as the operator has been recognized.

### 5.4 The standard library

> Name | Recognizes | Produces | Comment
> endrule(X) | X | α, if X : α | (inlined)
> midrule(X) | X | α, if X : α
> option(X) | є | X | α option, if X : α | (also X?)
> ioption(X) | є | X | α option, if X : α | (inlined)
> boption(X) | є | X | bool
> loption(X) | є | X | α list, if X : α list
> pair(X, Y) | X Y | α×β, if X : α and Y : β
> separated_pair(X, sep, Y) | X sep Y | α×β, if X : α and Y : β
> preceded(opening, X) | opening X | α, if X : α
> terminated(X, closing) | X closing | α, if X : α
> delimited(opening, X, closing) | opening X closing | α, if X : α
> list(X) | a possibly empty sequence of X’s | α list, if X : α | (also X*)
> nonempty_list(X) | a nonempty sequence of X’s | α list, if X : α | (also X+)
> separated_list(sep, X) | a possibly empty sequence of X’s separated with sep’s | α list, if X : α
> separated_nonempty_list(sep, X) | a nonempty sequence of X’s separated with sep’s | α list, if X : α
> rev(X) | X | α list, if X : α list | (inlined)
> flatten(X) | X | α list, if X : α list list | (inlined)
> append(X, Y) | X Y | α list, if X, Y : α list | (inlined)

> Figure 3: Summary of the standard library; see standard.mly for details

Once equipped with a rudimentary module system (§5.1), parameterization (§5.2), and inlining (§5.3), it is straightforward to propose a collection of commonly used definitions, such as options, sequences, lists, and so on. This standard library is joined, by default, with every grammar specification. A summary of the nonterminal symbols offered by the standard library appears in Figure 3. See also the short-hands documented in Figure 2.

By relying on the standard library, a client module can concisely define more elaborate notions. For instance, the following rule:

> %inline plist(X):
> | xs = loption(delimited(LPAREN, separated_nonempty_list(COMMA, X), RPAREN)) { xs }

causes plist(X) to recognize a list of X’s, where the empty list is represented by the empty string, and a non-empty list is delimited with parentheses and comma-separated.

The standard library is stored in a file named standard.mly, which is embedded inside Menhir when it is built. The command line switch --no-stdlib instructs Menhir to not load the standard library.

The meaning of the symbols defined in the standard library (Figure 3) should be clear in most cases. Yet, the symbols endrule(X) and midrule(X) deserve an explanation. Both take an argument X, which typically will be instantiated with an anonymous rule (§4.2.4). Both are defined as a synonym for X. In both cases, this allows placing an anonymous subrule in the middle of a rule.

For instance, the following is a well-formed production:

cat endrule(dog { OCaml code1 }) cow { OCaml code2 }

This production consists of three producers, namely cat and endrule(dog { OCaml code1 }) and cow, and a semantic action { OCaml code2 }. Because endrule(X) is declared as an %inline synonym for X, the expansion of anonymous rules (§4.2.4), followed with the expansion of %inline symbols (§5.3), transforms the above production into the following:

cat dog cow { OCaml code1; OCaml code2 }

Note that OCaml code1 moves to the end of the rule, which means that this code is executed only after cat, dog and cow have been recognized. In this example, the use of endrule is rather pointless, as the expanded code is more concise and clearer than the original code. Still, endrule can be useful when its actual argument is an anonymous rule with multiple branches.

midrule is used in exactly the same way as endrule, but its expansion is different. For instance, the following is a well-formed production:

cat midrule({ OCaml code1 }) cow { OCaml code2 }

(There is no dog in this example; this is intentional.) Because midrule(X) is a synonym for X, but is not declared %inline, the expansion of anonymous rules (§4.2.4), followed with the expansion of %inline symbols (§5.3), transforms the above production into the following:

cat xxx cow { OCaml code2 }

where the fresh nonterminal symbol xxx is separately defined by the rule xxx: { OCaml code1 } . Thus, xxx recognizes the empty string, and as soon as it is recognized, OCaml code1 is executed. This is known as a “mid-rule action”.

## 6 Conflicts

When a shift/reduce or reduce/reduce conflict is detected, it is classified as either benign, if it can be resolved by consulting user-supplied precedence declarations, or severe, if it cannot. Benign conflicts are not reported. Severe conflicts are reported and, if the --explain switch is on, explained.

### 6.1 When is a conflict benign?

A shift/reduce conflict involves a single token (the one that one might wish to shift) and one or more productions (those that one might wish to reduce). When such a conflict is detected, the precedence level (§4.1.4, §4.2.1) of these entities are looked up and compared as follows:

1. if only one production is involved, and if it has higher priority than the token, then the conflict is resolved in favor of reduction.
2. if only one production is involved, and if it has the same priority as the token, then the associativity status of the token is looked up:
   1. if the token was declared nonassociative, then the conflict is resolved in favor of neither action, that is, a syntax error will be signaled if this token shows up when this production is about to be reduced;
   2. if the token was declared left-associative, then the conflict is resolved in favor of reduction;
   3. if the token was declared right-associative, then the conflict is resolved in favor of shifting.
3. if multiple productions are involved, and if, considered one by one, they all cause the conflict to be resolved in the same way (that is, either in favor in shifting, or in favor of neither), then the conflict is resolved in that way.

In either of these cases, the conflict is considered benign. Otherwise, it is considered severe. Note that a reduce/reduce conflict is always considered severe, unless it happens to be subsumed by a benign multi-way shift/reduce conflict (item 3 above).

### 6.2 How are severe conflicts explained?

When the --dump switch is on, a description of the automaton is written to the .automaton file. Severe conflicts are shown as part of this description. Fortunately, there is also a way of understanding conflicts in terms of the grammar, rather than in terms of the automaton. When the --explain switch is on, a textual explanation is written to the .conflicts file.

Not all conflicts are explained in this file: instead, only one conflict per automaton state is explained. This is done partly in the interest of brevity, but also because Pager’s algorithm can create artificial conflicts in a state that already contains a true LR(1) conflict; thus, one cannot hope in general to explain all of the conflicts that appear in the automaton. As a result of this policy, once all conflicts explained in the .conflicts file have been fixed, one might need to run Menhir again to produce yet more conflict explanations.

> > %token IF THEN ELSE
> > %start < expression > expression
> > %%
> > expression:
> > | …
> > | IF b = expression THEN e = expression { … }
> > | IF b = expression THEN e = expression ELSE f = expression { … }
> > | …

> Figure 4: Basic example of a shift/reduce conflict

#### How the conflict state is reached

Figure 4 shows a grammar specification with a typical shift/reduce conflict. When this specification is analyzed, the conflict is detected, and an explanation is written to the .conflicts file. The explanation first indicates in which state the conflict lies by showing how that state is reached. Here, it is reached after recognizing the following string of terminal and nonterminal symbols—the conflict string:

> IF expression THEN IF expression THEN expression

Allowing the conflict string to contain both nonterminal and terminal symbols usually makes it shorter and more readable. If desired, a conflict string composed purely of terminal symbols could be obtained by replacing each occurrence of a nonterminal symbol N with an arbitrary N-sentence.

The conflict string can be thought of as a path that leads from one of the automaton’s start states to the conflict state. When multiple such paths exist, the one that is displayed is chosen shortest. Nevertheless, it may sometimes be quite long. In that case, artificially (and temporarily) declaring some existing nonterminal symbols to be start symbols has the effect of adding new start states to the automaton and can help produce shorter conflict strings. Here, expression was declared to be a start symbol, which is why the conflict string is quite short.

In addition to the conflict string, the .conflicts file also states that the conflict token is ELSE. That is, when the automaton has recognized the conflict string and when the lookahead token (the next token on the input stream) is ELSE, a conflict arises. A conflict corresponds to a choice: the automaton is faced with several possible actions, and does not know which one should be taken. This indicates that the grammar is not LR(1). The grammar may or may not be inherently ambiguous.

In our example, the conflict string and the conflict token are enough to understand why there is a conflict: when two IF constructs are nested, it is ambiguous which of the two constructs the ELSE branch should be associated with. Nevertheless, the .conflicts file provides further information: it explicitly shows that there exists a conflict, by proving that two distinct actions are possible. Here, one of these actions consists in shifting, while the other consists in reducing: this is a shift/reduce conflict.

A proof takes the form of a partial derivation tree whose fringe begins with the conflict string, followed by the conflict token. A derivation tree is a tree whose nodes are labeled with symbols. The root node carries a start symbol. A node that carries a terminal symbol is considered a leaf, and has no children. A node that carries a nonterminal symbol N either is considered a leaf, and has no children; or is not considered a leaf, and has n children, where n≥ 0, labeled x1,…,xn, where N → x1,…,xn is a production. The fringe of a partial derivation tree is the string of terminal and nonterminal symbols carried by the tree’s leaves. A string of terminal and nonterminal symbols that is the fringe of some partial derivation tree is a sentential form.

#### Why shifting is legal

> Figure 5: A partial derivation tree that justifies shifting

> | expression
> | IF expression THEN | expression
> |  | IF expression THEN expression . ELSE expression

> Figure 6: A textual version of the tree in Figure 5

In our example, the proof that shifting is possible is the derivation tree shown in Figures 5 and 6. At the root of the tree is the grammar’s start symbol, expression. This symbol develops into the string IF expression THEN expression, which forms the tree’s second level. The second occurrence of expression in that string develops into IF expression THEN expression ELSE expression, which forms the tree’s last level. The tree’s fringe, a sentential form, is the string IF expression THEN IF expression THEN expression ELSE expression. As announced earlier, it begins with the conflict string IF expression THEN IF expression THEN expression, followed with the conflict token ELSE.

In Figure 6, the end of the conflict string is materialized with a dot. Note that this dot does not occupy the rightmost position in the tree’s last level. In other words, the conflict token (ELSE) itself occurs on the tree’s last level. In practical terms, this means that, after the automaton has recognized the conflict string and peeked at the conflict token, it makes sense for it to shift that token.

#### Why reducing is legal

> Figure 7: A partial derivation tree that justifies reducing

> | expression
> | IF expression THEN | expression ELSE expression // lookahead token appears
> |  | IF expression THEN expression .

> Figure 8: A textual version of the tree in Figure 7

In our example, the proof that shifting is possible is the derivation tree shown in Figures 7 and 8. Again, the sentential form found at the fringe of the tree begins with the conflict string, followed with the conflict token.

Again, in Figure 8, the end of the conflict string is materialized with a dot. Note that, this time, the dot occupies the rightmost position in the tree’s last level. In other words, the conflict token (ELSE) appeared on an earlier level (here, on the second level). This fact is emphasized by the comment // lookahead token appears found at the second level. In practical terms, this means that, after the automaton has recognized the conflict string and peeked at the conflict token, it makes sense for it to reduce the production that corresponds to the tree’s last level—here, the production is expression → IF expression THEN expression.

#### An example of a more complex derivation tree

Figures 9 and 10 show a partial derivation tree that justifies reduction in a more complex situation. (This derivation tree is relative to a grammar that is not shown.) Here, the conflict string is DATA UIDENT EQUALS UIDENT; the conflict token is LIDENT. It is quite clear that the fringe of the tree begins with the conflict string. However, in this case, the fringe does not explicitly exhibit the conflict token. Let us examine the tree more closely and answer the question: following UIDENT, what’s the next terminal symbol on the fringe?

> Figure 9: A partial derivation tree that justifies reducing

> | decls
> | decl opt_semi decls // lookahead token appears because opt_semi can vanish and decls can begin with LIDENT
> | DATA UIDENT EQUALS | tycon_expr // lookahead token is inherited
> |  | tycon_item // lookahead token is inherited
> |  | UIDENT | opt_type_exprs // lookahead token is inherited
> |  |  | .

> Figure 10: A textual version of the tree in Figure 9

First, note that opt_type_exprs is not a leaf node, even though it has no children. The grammar contains the production opt_type_exprs → є: the nonterminal symbol opt_type_exprs develops to the empty string. (This is made clear in Figure 10, where a single dot appears immediately below opt_type_exprs.) Thus, opt_type_exprs is not part of the fringe.

Next, note that opt_type_exprs is the rightmost symbol within its level. Thus, in order to find the next symbol on the fringe, we have to look up one level. This is the meaning of the comment // lookahead token is inherited. Similarly, tycon_item and tycon_expr appear rightmost within their level, so we again have to look further up.

This brings us back to the tree’s second level. There, decl is not the rightmost symbol: next to it, we find opt_semi and decls. Does this mean that opt_semi is the next symbol on the fringe? Yes and no. opt_semi is a nonterminal symbol, but we are really interested in finding out what the next terminal symbol on the fringe could be. The partial derivation tree shown in Figures 9 and 10 does not explicitly answer this question. In order to answer it, we need to know more about opt_semi and decls.

Here, opt_semi stands (as one might have guessed) for an optional semicolon, so the grammar contains a production opt_semi → є. This is indicated by the comment // opt_semi can vanish. (Nonterminal symbols that generate є are also said to be nullable.) Thus, one could choose to turn this partial derivation tree into a larger one by developing opt_semi into є, making it a non-leaf node. That would yield a new partial derivation tree where the next symbol on the fringe, following UIDENT, is decls.

Now, what about decls? Again, it is a nonterminal symbol, and we are really interested in finding out what the next terminal symbol on the fringe could be. Again, we need to imagine how this partial derivation tree could be turned into a larger one by developing decls. Here, the grammar happens to contain a production of the form decls → LIDENT … This is indicated by the comment // decls can begin with LIDENT. Thus, by developing decls, it is possible to construct a partial derivation tree where the next symbol on the fringe, following UIDENT, is LIDENT. This is precisely the conflict token.

To sum up, there exists a partial derivation tree whose fringe begins with the conflict string, followed with the conflict token. Furthermore, in that derivation tree, the dot occupies the rightmost position in the last level. As in our previous example, this means that, after the automaton has recognized the conflict string and peeked at the conflict token, it makes sense for it to reduce the production that corresponds to the tree’s last level—here, the production is opt_type_exprs → є.

#### Greatest common factor among derivation trees

Understanding conflicts requires comparing two (or more) derivation trees. It is frequent for these trees to exhibit a common factor, that is, to exhibit identical structure near the top of the tree, and to differ only below a specific node. Manual identification of that node can be tedious, so Menhir performs this work automatically. When explaining a n-way conflict, it first displays the greatest common factor of the n derivation trees. A question mark symbol (?) is used to identify the node where the trees begin to differ. Then, Menhir displays each of the n derivation trees, without their common factor – that is, it displays n sub-trees that actually begin to differ at the root. This should make visual comparisons significantly easier.

### 6.3 How are severe conflicts resolved in the end?

It is unspecified how severe conflicts are resolved. Menhir attempts to mimic ocamlyacc’s specification, that is, to resolve shift/reduce conflicts in favor of shifting, and to resolve reduce/reduce conflicts in favor of the production that textually appears earliest in the grammar specification. However, this specification is inconsistent in case of three-way conflicts, that is, conflicts that simultaneously involve a shift action and several reduction actions. Furthermore, textual precedence can be undefined when the grammar specification is split over multiple modules. In short, Menhir’s philosophy is that

severe conflicts should not be tolerated,

so you should not care how they are resolved.

### 6.4 End-of-stream conflicts

Menhir’s treatment of the end of the token stream is (believed to be) fully compatible with ocamlyacc’s. Yet, Menhir attempts to be more user-friendly by warning about a class of so-called “end-of-stream conflicts”.

#### How the end of stream is handled

In many textbooks on parsing, it is assumed that the lexical analyzer, which produces the token stream, produces a special token, written #, to signal that the end of the token stream has been reached. A parser generator can take advantage of this by transforming the grammar: for each start symbol S in the original grammar, a new start symbol S’ is defined, together with the production S′→ S# . The symbol S is no longer a start symbol in the new grammar. This means that the parser will accept a sentence derived from S only if it is immediately followed by the end of the token stream.

This approach has the advantage of simplicity. However, ocamlyacc and Menhir do not follow it, for several reasons. Perhaps the most convincing one is that it is not flexible enough: sometimes, it is desirable to recognize a sentence derived from S, without requiring that it be followed by the end of the token stream: this is the case, for instance, when reading commands, one by one, on the standard input channel. In that case, there is no end of stream: the token stream is conceptually infinite. Furthermore, after a command has been recognized, we do not wish to examine the next token, because doing so might cause the program to block, waiting for more input.

In short, ocamlyacc and Menhir’s approach is to recognize a sentence derived from S and to not look, if possible, at what follows. However, this is possible only if the definition of S is such that the end of an S-sentence is identifiable without knowledge of the lookahead token. When the definition of S does not satisfy this criterion, and end-of-stream conflict arises: after a potential S-sentence has been read, there can be a tension between consulting the next token, in order to determine whether the sentence is continued, and not consulting the next token, because the sentence might be over and whatever follows should not be read. Menhir warns about end-of-stream conflicts, whereas ocamlyacc does not.

#### A definition of end-of-stream conflicts

Technically, Menhir proceeds as follows. A # symbol is introduced. It is, however, only a pseudo-token: it is never produced by the lexical analyzer. For each start symbol S in the original grammar, a new start symbol S’ is defined, together with the production S′→ S. The corresponding start state of the LR(1) automaton is composed of the LR(1) item S′ → . S [# ]. That is, the pseudo-token # initially appears in the lookahead set, indicating that we expect to be done after recognizing an S-sentence. During the construction of the LR(1) automaton, this lookahead set is inherited by other items, with the effect that, in the end, the automaton has:

- shift actions only on physical tokens; and
- reduce actions either on physical tokens or on the pseudo-token #.

A state of the automaton has a reduce action on # if, in that state, an S-sentence has been read, so that the job is potentially finished. A state has a shift or reduce action on a physical token if, in that state, more tokens potentially need to be read before an S-sentence is recognized. If a state has a reduce action on #, then that action should be taken without requesting the next token from the lexical analyzer. On the other hand, if a state has a shift or reduce action on a physical token, then the lookahead token must be consulted in order to determine if that action should be taken.

> > %token < int > INT
> > %token PLUS TIMES
> > %left PLUS
> > %left TIMES
> > %start < int > expr
> > %%
> > expr:
> > | i = INT { i }
> > | e1 = expr PLUS e2 = expr { e1 + e2 }
> > | e1 = expr TIMES e2 = expr { e1 * e2 }

> Figure 11: Basic example of an end-of-stream conflict

>     State 6:
>     expr -> expr . PLUS expr [ # TIMES PLUS ]
>     expr -> expr PLUS expr . [ # TIMES PLUS ]
>     expr -> expr . TIMES expr [ # TIMES PLUS ]
>     -- On TIMES shift to state 3
>     -- On # PLUS reduce production expr -> expr PLUS expr

>     State 4:
>     expr -> expr . PLUS expr [ # TIMES PLUS ]
>     expr -> expr . TIMES expr [ # TIMES PLUS ]
>     expr -> expr TIMES expr . [ # TIMES PLUS ]
>     -- On # TIMES PLUS reduce production expr -> expr TIMES expr

>     State 2:
>     expr' -> expr . [ # ]
>     expr -> expr . PLUS expr [ # TIMES PLUS ]
>     expr -> expr . TIMES expr [ # TIMES PLUS ]
>     -- On TIMES shift to state 3
>     -- On PLUS shift to state 5
>     -- On # accept expr

> Figure 12: Part of an LR automaton for the grammar in Figure 11

> > …
> > %token END
> > %start < int > main // instead of expr
> > %%
> > main:
> > | e = expr END { e }
> > expr:
> > | …

> Figure 13: Fixing the grammar specification in Figure 11

An end-of-stream conflict arises when a state has distinct actions on # and on at least one physical token. In short, this means that the end of an S-sentence cannot be unambiguously identified without examining one extra token. Menhir’s default behavior, in that case, is to suppress the action on #, so that more input is always requested.

#### Example

Figure 11 shows a grammar that has end-of-stream conflicts. When this grammar is processed, Menhir warns about these conflicts, and further warns that expr is never accepted. Let us explain.

Part of the corresponding automaton, as described in the .automaton file, is shown in Figure 12. Explanations at the end of the .automaton file (not shown) point out that states 6 and 2 have an end-of-stream conflict. Indeed, both states have distinct actions on # and on the physical token TIMES. It is interesting to note that, even though state 4 has actions on # and on physical tokens, it does not have an end-of-stream conflict. This is because the action taken in state 4 is always to reduce the production expr → expr TIMES expr, regardless of the lookahead token.

By default, Menhir produces a parser where end-of-stream conflicts are resolved in favor of looking ahead: that is, the problematic reduce actions on # are suppressed. This means, in particular, that the accept action in state 2, which corresponds to reducing the production expr → expr’, is suppressed. This explains why the symbol expr is never accepted: because expressions do not have an unambiguous end marker, the parser will always request one more token and will never stop.

In order to avoid this end-of-stream conflict, the standard solution is to introduce a new token, say END, and to use it as an end marker for expressions. The END token could be generated by the lexical analyzer when it encounters the actual end of stream, or it could correspond to a piece of concrete syntax, say, a line feed character, a semicolon, or an end keyword. The solution is shown in Figure 13.

## 7 Positions

When an ocamllex-generated lexical analyzer produces a token, it updates two fields, named `lex_start_p` and `lex_curr_p`, in its environment record, whose type is `Lexing.lexbuf`. Each of these fields holds a value of type `Lexing.position`. Together, they represent the token’s start and end positions within the text that is being scanned. These fields are read by Menhir after calling the lexical analyzer, so it is the lexical analyzer’s responsibility to correctly set these fields.

A position consists mainly of an offset (the position’s `pos_cnum` field), but also holds information about the current file name, the current line number, and the current offset within the current line. (Not all ocamllex-generated analyzers keep this extra information up to date. This must be explicitly programmed by the author of the lexical analyzer.)

> `$startpos` |  | start position of the first symbol in the production’s right-hand side, if there is one;
> |  | end position of the most recently parsed symbol, otherwise
> `$endpos` |  | end position of the last symbol in the production’s right-hand side, if there is one;
> |  | end position of the most recently parsed symbol, otherwise
> `$startpos(` `$`i | id `)` |  | start position of the symbol named `$`i or id
> `$endpos(` `$`i | id `)` |  | end position of the symbol named `$`i or id
> $symbolstartpos |  | start position of the leftmost symbol id such that `$startpos(`id`)` `!=` `$endpos(`id`)`;
> |  | if there is no such symbol, `$endpos`
> `$startofs` |
> `$endofs` |
> `$startofs(` `$`i | id `)` |  | same as above, but produce an integer offset instead of a position
> `$endofs(` `$`i | id `)` |
> `$symbolstartofs` |
> `$loc` |  | stands for the pair `($startpos, $endpos)`
> `$loc(` id `)` |  | stands for the pair `($startpos(` id `), $endpos(` id `))`
> `$sloc` |  | stands for the pair `($symbolstartpos, $endpos)`

> Figure 14: Position-related keywords

> `symbol_start_pos()` | $symbolstartpos |
> `symbol_end_pos()` | `$endpos` |
> `rhs_start_pos i` | `$startpos($i)` |  | (1 ≤ i ≤ n)
> `rhs_end_pos i` | `$endpos($i)` |  | (1 ≤ i ≤ n)
> `symbol_start()` | `$symbolstartofs` |
> `symbol_end()` | `$endofs` |
> `rhs_start i` | `$startofs($i)` |  | (1 ≤ i ≤ n)
> `rhs_end i` | `$endofs($i)` |  | (1 ≤ i ≤ n)

> Figure 15: Translating position-related incantations from ocamlyacc to Menhir

This mechanism allows associating pairs of positions with terminal symbols. If desired, Menhir automatically extends it to nonterminal symbols as well. That is, it offers a mechanism for associating pairs of positions with terminal or nonterminal symbols. This is done by making a set of keywords available to semantic actions (Figure 14). These keywords are not available outside of a semantic action: in particular, they cannot be used within an OCaml header.

OCaml’s standard library module Parsing is deprecated. The functions that it offers can be called, but will return dummy positions.

We remark that, if the current production has an empty right-hand side, then `$startpos` and `$endpos` are equal, and (by convention) are the end position of the most recently parsed symbol (that is, the symbol that happens to be on top of the automaton’s stack when this production is reduced). If the current production has a nonempty right-hand side, then `$startpos` is the same as `$startpos($1)` and `$endpos` is the same as `$endpos($`n`)`, where n is the length of the right-hand side.

More generally, if the current production has matched a sentence of length zero, then `$startpos` and `$endpos` will be equal, and conversely.

The position `$startpos` is sometimes “further towards the left” than one would like. For example, in the following production:

      declaration: modifier? variable { $startpos }

the keyword `$startpos` represents the start position of the optional modifier `modifier?`. If this modifier turns out to be absent, then its start position is (by definition) the end position of the most recently parsed symbol. This may not be what is desired: perhaps the user would prefer in this case to use the start position of the symbol `variable`. This is achieved by using $symbolstartpos instead of `$startpos`. By definition, $symbolstartpos is the start position of the leftmost symbol whose start and end positions differ. In this example, the computation of $symbolstartpos skips the absent `modifier`, whose start and end positions coincide, and returns the start position of the symbol `variable` (assuming this symbol has distinct start and end positions).

There is no keyword `$symbolendpos`. Indeed, the problem with `$startpos` is due to the asymmetry in the definition of `$startpos` and `$endpos` in the case of an empty right-hand side, and does not affect `$endpos`.

The positions computed by Menhir are exactly the same as those computed by `ocamlyacc`1. More precisely, Figure 15 sums up how to translate a call to the Parsing module, as used in an ocamlyacc grammar, to a Menhir keyword.

We note that Menhir’s `$startpos` does not appear in the right-hand column in Figure 15. In other words, Menhir’s `$startpos` does not correspond exactly to any of the ocamlyacc function calls. An exact ocamlyacc equivalent of `$startpos` is `rhs_start_pos 1` if the current production has a nonempty right-hand side and `symbol_start_pos()` if it has an empty right-hand side.

Finally, we remark that Menhir’s %inline keyword (§5.3) does not affect the computation of positions. The same positions are computed, regardless of where %inline keywords are placed.

## 8 Using Menhir as an interpreter

When --interpret is set, Menhir no longer behaves as a compiler. Instead, it acts as an interpreter. That is, it repeatedly:

- reads a sentence off the standard input channel;
- parses this sentence, according to the grammar;
- displays an outcome.

This process stops when the end of the input channel is reached.

### 8.1 Sentences

The syntax of sentences is as follows:

sentence | ::= | [ lid : ] uid … uid \n

Less formally, a sentence is a sequence of zero or more terminal symbols (uid’s), separated with whitespace, terminated with a newline character, and optionally preceded with a non-terminal start symbol (lid). This non-terminal symbol can be omitted if, and only if, the grammar only has one start symbol.

For instance, here are four valid sentences for the grammar of arithmetic expressions found in the directory demos/calc:

    main: INT PLUS INT EOL
    INT PLUS INT
    INT PLUS PLUS INT EOL
    INT PLUS PLUS

In the first sentence, the start symbol main was explicitly specified. In the other sentences, it was omitted, which is permitted, because this grammar has no start symbol other than main. The first sentence is a stream of four terminal symbols, namely INT, PLUS, INT, and EOL. These terminal symbols must be provided under their symbolic names. Writing, say, “12+32\n” instead of INT PLUS INT EOL is not permitted. Menhir would not be able to make sense of such a concrete notation, since it does not have a lexer for it.

### 8.2 Outcomes

As soon as Menhir is able to read a complete sentence off the standard input channel (that is, as soon as it finds the newline character that ends the sentence), it parses the sentence according to whichever grammar was specified on the command line, and displays an outcome.

An outcome is one of the following:

- ACCEPT: a prefix of the sentence was successfully parsed; a parser generated by Menhir would successfully stop and produce a semantic value;
- OVERSHOOT: the end of the sentence was reached before it could be accepted; a parser generated by Menhir would request a non-existent “next token” from the lexer, causing it to fail or block;
- REJECT: the sentence was not accepted; a parser generated by Menhir would raise the exception Error.

When --interpret-show-cst is set, each ACCEPT outcome is followed with a concrete syntax tree. A concrete syntax tree is either a leaf or a node. A leaf is either a terminal symbol or error. A node is annotated with a non-terminal symbol, and carries a sequence of immediate descendants that correspond to a valid expansion of this non-terminal symbol. Menhir’s notation for concrete syntax trees is as follows:

cst | ::= | uid
|  | error
|  | [ lid : cst … cst ]

For instance, if one wished to parse the example sentences of §8.1 using the grammar of arithmetic expressions in demos/calc, one could invoke Menhir as follows:

    $ menhir --interpret --interpret-show-cst demos/calc/parser.mly
    main: INT PLUS INT EOL
    ACCEPT
    [main: [expr: [expr: INT] PLUS [expr: INT]] EOL]
    INT PLUS INT
    OVERSHOOT
    INT PLUS PLUS INT EOL
    REJECT
    INT PLUS PLUS
    REJECT

(Here, Menhir’s input—the sentences provided by the user on the standard input channel— is shown intermixed with Menhir’s output—the outcomes printed by Menhir on the standard output channel.) The first sentence is valid, and accepted; a concrete syntax tree is displayed. The second sentence is incomplete, because the grammar specifies that a valid expansion of main ends with the terminal symbol EOL; hence, the outcome is OVERSHOOT. The third sentence is invalid, because of the repeated occurrence of the terminal symbol PLUS; the outcome is REJECT. The fourth sentence, a prefix of the third one, is rejected for the same reason.

### 8.3 Remarks

Using Menhir as an interpreter offers an easy way of debugging your grammar. For instance, if one wished to check that addition is considered left-associative, as requested by the %left directive found in the file demos/calc/parser.mly, one could submit the following sentence:

    $ ./menhir --interpret --interpret-show-cst ../demos/calc/parser.mly
    INT PLUS INT PLUS INT EOL
    ACCEPT
    [main:
      [expr: [expr: [expr: INT] PLUS [expr: INT]] PLUS [expr: INT]]
      EOL
    ]

The concrete syntax tree displayed by Menhir is skewed towards the left, as desired.

The switches --interpret and --trace can be used in conjunction. When --trace is set, the interpreter logs its actions to the standard error channel.

## 9 Generated API

When Menhir processes a grammar specification, say parser.mly, it produces one OCaml module, Parser, whose code resides in the file parser.ml and whose signature resides in the file parser.mli. We now review this signature. For simplicity, we assume that the grammar specification has just one start symbol `main`, whose OCaml type is `thing`.

### 9.1 Monolithic API

The monolithic API defines the type `token`, the exception `Error`, and the parsing function `main`, named after the start symbol of the grammar.

The type `token` is an algebraic data type. A value of type `token` represents a terminal symbol and its semantic value. For instance, if the grammar contains the declarations `%token A` and `%token<int> B`, then the generated file parser.mli contains the following definition:

      type token =
      | A
      | B of int

If --only-tokens is specified on the command line, the type `token` is generated, and the rest is omitted. On the contrary, if --external-tokens is used, the type `token` is omitted, but the rest (described below) is generated.

The exception `Error` carries no argument. It is raised by the parsing function `main` (described below) when a syntax error is detected.

      exception Error

Next comes one parsing function for each start symbol of the grammar. Here, we have assumed that there is one start symbol, named `main`, so the generated file parser.mli contains the following declaration:

      val main: (Lexing.lexbuf -> token) -> Lexing.lexbuf -> thing

This function expects two arguments, namely: a lexer, which typically is produced by ocamllex and has type `Lexing.lexbuf -> token`; and a lexing buffer, which has type `Lexing.lexbuf`. This API is compatible with ocamlyacc. (For information on using Menhir without ocamllex, please consult §16.) This API is “monolithic” in the sense that there is just one function, which does everything: it pulls tokens from the lexer, parses, and eventually returns a semantic value (or fails by throwing the exception Error).

### 9.2 Incremental API

If --table is set, Menhir offers an incremental API in addition to the monolithic API. In this API, control is inverted. The parser does not have access to the lexer. Instead, when the parser needs the next token, it stops and returns its current state to the user. The user is then responsible for obtaining this token (typically by invoking the lexer) and resuming the parser from that state. The directory demos/calc-incremental contains a demo that illustrates the use of the incremental API.

This API is “incremental” in the sense that the user has access to a sequence of the intermediate states of the parser. Assuming that semantic values are immutable, a parser state is a persistent data structure: it can be stored and used multiple times, if desired. This enables applications such as “live parsing”, where a buffer is continuously parsed while it is being edited. The parser can be re-started in the middle of the buffer whenever the user edits a character. Because two successive parser states share most of their data in memory, a list of n successive parser states occupies only O(n) space in memory.

#### 9.2.1 Starting the parser

In this API, the parser is started by invoking `Incremental.main`. (Recall that we assume that `main` is the name of the start symbol.) The generated file parser.mli contains the following declaration:

      module Incremental : sig
        val main: position -> thing MenhirInterpreter.checkpoint
      end

The argument is the initial position. If the lexer is based on an OCaml lexing buffer, this argument should be `lexbuf.lex_curr_p`. In §9.2 and §9.3, the type `position` is a synonym for `Lexing.position`.

We emphasize that the function `Incremental.main` does not parse anything. It constructs a checkpoint which serves as a starting point. The functions `offer` and `resume`, described below, are used to drive the parser.

#### 9.2.2 Driving the parser

The sub-module MenhirInterpreter is also part of the incremental API. Its declaration, which appears in the generated file parser.mli, is as follows:

      module MenhirInterpreter : MenhirLib.IncrementalEngine.INCREMENTAL_ENGINE
        with type token = token

The signature `INCREMENTAL_ENGINE`, defined in the module MenhirLib.IncrementalEngine, contains many types and functions, which are described in the rest of this section (§9.2.2) and in the following sections (§9.2.3, §9.2.4).

Please keep in mind that, from the outside, these types and functions should be referred to with an appropriate prefix. For instance, the type `checkpoint` should be referred to as `MenhirInterpreter.checkpoint`, or `Parser.MenhirInterpreter.checkpoint`, depending on which modules the user chooses to open.

      type 'a env

The abstract type `'a env` represents the current state of the parser. (That is, it contains the current state and stack of the LR automaton.) Assuming that semantic values are immutable, it is a persistent data structure: it can be stored and used multiple times, if desired. The parameter `'a` is the type of the semantic value that will eventually be produced if the parser succeeds.

      type production

The abstract type `production` represents a production of the grammar. The “start productions” (which do not exist in an .mly file, but are constructed by Menhir internally) are not part of this type.

      type 'a checkpoint = private
        | InputNeeded of 'a env
        | Shifting of 'a env * 'a env * bool
        | AboutToReduce of 'a env * production
        | HandlingError of 'a env
        | Accepted of 'a
        | Rejected

The type `'a checkpoint` represents an intermediate or final state of the parser. An intermediate checkpoint is a suspension: it records the parser’s current state, and allows parsing to be resumed. The parameter `'a` is the type of the semantic value that will eventually be produced if the parser succeeds.

`Accepted` and `Rejected` are final checkpoints. `Accepted` carries a semantic value.

`InputNeeded` is an intermediate checkpoint. It means that the parser wishes to read one token before continuing.

`Shifting` is an intermediate checkpoint. It means that the parser is taking a shift transition. It exposes the state of the parser before and after the transition. The Boolean parameter tells whether the parser intends to request a new token after this transition. (It always does, except when it is about to accept.)

`AboutToReduce` is an intermediate checkpoint: it means that the parser is about to perform a reduction step. `HandlingError` is also an intermediate checkpoint: it means that the parser has detected an error and is about to handle it. (Error handling is typically performed in several steps, so the next checkpoint is likely to be `HandlingError` again.) In these two cases, the parser does not need more input. The parser suspends itself at this point only in order to give the user an opportunity to observe the parser’s transitions and possibly handle errors in a different manner, if desired.

      val offer:
        'a checkpoint ->
        token * position * position ->
        'a checkpoint

The function `offer` allows the user to resume the parser after the parser has suspended itself with a checkpoint of the form `InputNeeded env`. This function expects the previous checkpoint `checkpoint` as well as a new token (together with the start and end positions of this token). It produces a new checkpoint, which again can be an intermediate checkpoint or a final checkpoint. It does not raise any exception. (The exception Error is used only in the monolithic API.)

      val resume:
        'a checkpoint ->
        'a checkpoint

The function `resume` allows the user to resume the parser after the parser has suspended itself with a checkpoint of the form `AboutToReduce (env, prod)` or `HandlingError env`. This function expects just the previous checkpoint `checkpoint`. It produces a new checkpoint. It does not raise any exception.

The incremental API subsumes the monolithic API. Indeed, `main` can be (and is in fact) implemented by first using `Incremental.main`, then calling `offer` and `resume` in a loop, until a final checkpoint is obtained.

      type supplier =
        unit -> token * position * position

A token supplier is a function of no arguments which delivers a new token (together with its start and end positions) every time it is called. The function `loop` and its variants, described below, expect a supplier as an argument.

      val lexer_lexbuf_to_supplier:
        (Lexing.lexbuf -> token) -> Lexing.lexbuf -> supplier

The function `lexer_lexbuf_to_supplier`, applied to a lexer and to a lexing buffer, produces a fresh supplier.

The functions `offer` and `resume`, documented above, are sufficient to write a parser loop. One can imagine many variations of such a loop, which is why we expose `offer` and `resume` in the first place. Nevertheless, some variations are so common that it is worth providing them, ready for use. The following functions are implemented on top of `offer` and `resume`.

      val loop: supplier -> 'a checkpoint -> 'a

`loop supplier checkpoint` begins parsing from `checkpoint`, reading tokens from `supplier`. It continues parsing until it reaches a checkpoint of the form `Accepted v` or `Rejected`. In the former case, it returns `v`. In the latter case, it raises the exception `Error`. (By the way, this is how we implement the monolithic API on top of the incremental API.)

      val loop_handle:
        ('a -> 'answer) ->
        ('a checkpoint -> 'answer) ->
        supplier -> 'a checkpoint -> 'answer

`loop_handle succeed fail supplier checkpoint` begins parsing from `checkpoint`, reading tokens from `supplier`. It continues until it reaches a checkpoint of the form `Accepted v` or `HandlingError _` (or `Rejected`, but that should not happen, as `HandlingError _` will be observed first). In the former case, it calls `succeed v`. In the latter case, it calls `fail` with this checkpoint. It cannot raise `Error`.

This means that Menhir’s traditional error-handling procedure (which pops the stack until a state that can act on the error token is found) does not get a chance to run. Instead, the user can implement her own error handling code, in the `fail` continuation.

      val loop_handle_undo:
        ('a -> 'answer) ->
        ('a checkpoint -> 'a checkpoint -> 'answer) ->
        supplier -> 'a checkpoint -> 'answer

`loop_handle_undo` is analogous to `loop_handle`, but passes a pair of checkpoints (instead of a single checkpoint) to the failure continuation. The first (and oldest) checkpoint that is passed to the failure continuation is the last `InputNeeded` checkpoint that was encountered before the error was detected. The second (and newest) checkpoint is where the error was detected. (This is the same checkpoint that `loop_handle` would pass to its failure continuation.) Going back to the first checkpoint can be thought of as undoing any reductions that were performed after seeing the problematic token. (These reductions must be default reductions or spurious reductions.) This can be useful to someone who wishes to implement an error explanation or error recovery mechanism.

`loop_handle_undo` must be applied to an `InputNeeded` checkpoint. The initial checkpoint produced by `Incremental.main` is of this form.

      val shifts: 'a checkpoint -> 'a env option

`shifts checkpoint` assumes that `checkpoint` has been obtained by submitting a token to the parser. It runs the parser from `checkpoint`, through an arbitrary number of reductions, until the parser either accepts this token (i.e., shifts) or rejects it (i.e., signals an error). If the parser decides to shift, then `Some env` is returned, where `env` is the parser’s state just before shifting. Otherwise, `None` is returned. This can be used to test whether the parser is willing to accept a certain token. This function should be used with caution, though, as it causes semantic actions to be executed. It is desirable that all semantic actions be side-effect-free, or that their side-effects be harmless.

      val acceptable: 'a checkpoint -> token -> position -> bool

`acceptable checkpoint token pos` requires `checkpoint` to be an `InputNeeded` checkpoint. It returns `true` iff the parser is willing to shift this token. This can be used to test, after an error has been detected, which tokens would have been accepted at this point. To do this, one would typically use `loop_handle_undo` to get access to the last `InputNeeded` checkpoint that was encountered before the error was detected, and apply `acceptable` to that checkpoint.

`acceptable` is implemented using `shifts`, so, like `shifts`, it causes certain semantic actions to be executed. It is desirable that all semantic actions be side-effect-free, or that their side-effects be harmless.

#### 9.2.3 Inspecting the parser’s state

Although the type `env` is opaque, a parser state can be inspected via a few accessor functions, which are described in this section. The following types and functions are contained in the `MenhirInterpreter` sub-module.

      type 'a lr1state

The abstract type `'a lr1state` describes a (non-initial) state of the LR(1) automaton. If `s` is such a state, then `s` should have at least one incoming transition, and all of its incoming transitions carry the same (terminal or non-terminal) symbol, say A. We say that A is the incoming symbol of the state `s`. The index `'a` is the type of the semantic values associated with A. The role played by `'a` is clarified in the definition of the type `element`, which appears further on.

      val number: _ lr1state -> int

The states of the LR(1) automaton are numbered (from 0 and up). The function `number` maps a state to its number.

      val production_index: production -> int
      val find_production: int -> production

Productions are numbered. (The set of indices of all productions forms an interval, which does not necessarily begin at 0.) The function `production_index` converts a production to an integer number, whereas the function `find_production` carries out the reverse conversion. It is an error to apply `find_production` to an invalid index.

      type element =
        | Element: 'a lr1state * 'a * position * position -> element

The type `element` describes one entry in the stack of the LR(1) automaton. In a stack element of the form `Element (s, v, startp, endp)`, `s` is a (non-initial) state and `v` is a semantic value. The value `v` is associated with the incoming symbol A of the state `s`. In other words, the value `v` was pushed onto the stack just before the state `s` was entered. Thus, for some type `'a`, the state `s` has type `'a lr1state` and the value `v` has type `'a`. The positions `startp` and `endp` delimit the fragment of the input text that was reduced to the symbol A.

In order to do anything useful with the value `v`, one must gain information about the type `'a`, by inspection of the state `s`. So far, the type `'a lr1state` is abstract, so there is no way of inspecting `s`. The inspection API (§9.3) offers further tools for this purpose.

      val top: 'a env -> element option

`top env` returns the parser’s top stack element. The state contained in this stack element is the current state of the automaton. If the stack is empty, `None` is returned. In that case, the current state of the automaton must be an initial state.

      val pop_many: int -> 'a env -> 'a env option

`pop_many i env` pops `i` elements off the automaton’s stack. This is done via `i` successive invocations of `pop`. Thus, `pop_many 1` is `pop`. The index `i` must be nonnegative. The time complexity is O(i).

      val get: int -> 'a env -> element option

`get i env` returns the parser’s `i`-th stack element. The index `i` is 0-based: thus, `get 0` is `top`. If `i` is greater than or equal to the number of elements in the stack, `None` is returned. `get` is implemented using `pop_many` and `top`: its time complexity is O(i).

      val current_state_number: 'a env -> int

`current_state_number env` is the integer number of the automaton’s current state. Although this number might conceivably be obtained via the functions `top` and `number`, using `current_state_number` is preferable, because this method works even when the automaton’s stack is empty (in which case the current state is an initial state, and `top` returns `None`). This number can be passed as an argument to a `message` function generated by `menhir --compile-errors`.

      val equal: 'a env -> 'a env -> bool

`equal env1 env2` tells whether the parser configurations `env1` and `env2` are equal in the sense that the automaton’s current state is the same in `env1` and `env2` and the stack is physically the same in `env1` and `env2`. If `equal env1 env2` is `true`, then the sequence of the stack elements, as observed via `pop` and `top`, must be the same in `env1` and `env2`. Also, if `equal env1 env2` holds, then the checkpoints `input_needed env1` and `input_needed env2` must be equivalent. (The function `input_needed` is documented in §9.2.4.) The function `equal` has time complexity O(1).

      val positions: 'a env -> position * position

The function `positions` returns the start and end positions of the current lookahead token. If invoked in an initial state, this function returns a pair of twice the initial position that was passed as an argument to `main`.

      val env_has_default_reduction: 'a env -> bool
      val state_has_default_reduction: _ lr1state -> bool

When applied to an environment `env` taken from a checkpoint of the form `AboutToReduce (env, prod)`, the function `env_has_default_reduction` tells whether the reduction that is about to take place is a default reduction.

`state_has_default_reduction s` tells whether the state `s` has a default reduction. This includes the case where `s` is an accepting state.

#### 9.2.4 Updating the parser’s state

The functions presented in the previous section (§9.2.3) allow inspecting parser states of type `'a checkpoint` and `'a env`. However, so far, there are no functions for manufacturing new parser states, except `offer` and `resume`, which create new checkpoints by feeding tokens, one by one, to the parser.

In this section, a small number of functions are provided for manufacturing new parser states of type `'a env` and `'a checkpoint`. These functions allow going far back into the past and jumping ahead into the future, so to speak. In other words, they allow driving the parser in other ways than by feeding tokens into it. The functions `pop`, `force_reduction` and `feed` (part of the inspection API; see §9.3) construct values of type `'a env`. The function `input_needed` constructs values of type `'a checkpoint` and thereby allows resuming parsing in normal mode (via `offer`). Together, these functions can be used to implement error handling and error recovery strategies.

      val pop: 'a env -> 'a env option

`pop env` returns a new environment, where the parser’s top stack cell has been popped off. (If the stack is empty, `None` is returned.) This amounts to pretending that the (terminal or nonterminal) symbol that corresponds to this stack cell has not been read.

      val force_reduction: production -> 'a env -> 'a env

`force_reduction prod env` can be called only if in the state `env` the parser is capable of reducing the production `prod`. If this condition is satisfied, then this production is reduced, which means that its semantic action is executed (this can have side effects!) and the automaton makes a goto (nonterminal) transition. If this condition is not satisfied, an `Invalid_argument` exception is raised.

      val input_needed: 'a env -> 'a checkpoint

`input_needed env` returns `InputNeeded env`. Thus, out of a parser state that might have been obtained via a series of calls to the functions `pop`, `force_reduction`, `feed`, and so on, it produces a checkpoint, which can be used to resume normal parsing, by supplying this checkpoint as an argument to `offer`.

This function should be used with some care. It could “mess up the lookahead” in the sense that it allows parsing to resume in an arbitrary state `s` with an arbitrary lookahead symbol `t`, even though Menhir’s reachability analysis (which is carried out via the --list-errors switch) might well think that it is impossible to reach this particular configuration. If one is using Menhir’s new error reporting facility (§11), this could cause the parser to reach an error state for which no error message has been prepared.

### 9.3 Inspection API

If --inspection is set, Menhir offers an inspection API in addition to the monolithic and incremental APIs. (The reason why this is not done by default is that this requires more tables to be generated, thus making the generated parser larger.) Like the incremental API, the inspection API is found in the sub-module MenhirInterpreter. It offers the following types and functions.

The type `'a terminal` is a generalized algebraic data type (GADT). A value of type `'a terminal` represents a terminal symbol (without a semantic value). The index `'a` is the type of the semantic values associated with this symbol. For instance, if the grammar contains the declarations `%token A` and `%token<int> B`, then the generated module MenhirInterpreter contains the following definition:

      type _ terminal =
      | T_A : unit terminal
      | T_B : int terminal

The data constructors are named after the terminal symbols, prefixed with “`T_`”.

The type `'a nonterminal` is also a GADT. A value of type `'a nonterminal` represents a nonterminal symbol (without a semantic value). The index `'a` is the type of the semantic values associated with this symbol. For instance, if `main` is the only nonterminal symbol, then the generated module MenhirInterpreter contains the following definition:

      type _ nonterminal =
      | N_main : thing nonterminal

The data constructors are named after the nonterminal symbols, prefixed with “`N_`”.

The type `'a symbol` is the disjoint union of the types `'a terminal` and `'a nonterminal`. In other words, a value of type `'a symbol` represents a terminal or nonterminal symbol (without a semantic value). This type is (always) defined as follows:

      type 'a symbol =
        | T : 'a terminal -> 'a symbol
        | N : 'a nonterminal -> 'a symbol

The type `xsymbol` is an existentially quantified version of the type `'a symbol`. It is useful in situations where the index `'a` is not statically known. It is (always) defined as follows:

      type xsymbol =
        | X : 'a symbol -> xsymbol

The type `item` describes an LR(0) item, that is, a pair of a production `prod` and an index `i` into the right-hand side of this production. If the length of the right-hand side is `n`, then `i` is comprised between 0 and `n`, inclusive.

      type item =
          production * int

The following functions implement total orderings on the types `_ terminal`, `_ nonterminal`, `xsymbol`, `production`, and `item`.

      val compare_terminals: _ terminal -> _ terminal -> int
      val compare_nonterminals: _ nonterminal -> _ nonterminal -> int
      val compare_symbols: xsymbol -> xsymbol -> int
      val compare_productions: production -> production -> int
      val compare_items: item -> item -> int

The function `incoming_symbol` maps a (non-initial) LR(1) state `s` to its incoming symbol, that is, the symbol that the parser must recognize before it enters the state `s`.

      val incoming_symbol: 'a lr1state -> 'a symbol

This function can be used to gain access to the semantic value `v` in a stack element `Element (s, v, _, _)`. Indeed, by case analysis on the symbol `incoming_symbol s`, one gains information about the type `'a`, hence one obtains the ability to do something useful with the value `v`.

The function `items` maps a (non-initial) LR(1) state `s` to its LR(0) core, that is, to the underlying set of LR(0) items. This set is represented as a list, whose elements appear in an arbitrary order. This set is not closed under є-transitions.

      val items: _ lr1state -> item list

The functions `lhs` and `rhs` map a production `prod` to its left-hand side and right-hand side, respectively. The left-hand side is always a nonterminal symbol, hence always of the form `N _`. The right-hand side is a (possibly empty) sequence of (terminal or nonterminal) symbols.

      val lhs: production -> xsymbol
      val rhs: production -> xsymbol list

The function `nullable`, applied to a non-terminal symbol, tells whether this symbol is nullable. A nonterminal symbol is nullable if and only if it produces the empty word є.

      val nullable: _ nonterminal -> bool

The function call `first nt t` tells whether the FIRST set of the nonterminal symbol `nt` contains the terminal symbol `t`. That is, it returns `true` if and only if `nt` produces a word that begins with `t`. The function `xfirst` is identical to `first`, except it expects a first argument of type `xsymbol` instead of `_ terminal`.

      val first: _ nonterminal -> _ terminal -> bool
      val xfirst: xsymbol -> _ terminal -> bool

The function `foreach_terminal` enumerates the terminal symbols, including the special symbol error. The function `foreach_terminal_but_error` enumerates the terminal symbols, excluding error.

      val foreach_terminal:           (xsymbol -> 'a -> 'a) -> 'a -> 'a
      val foreach_terminal_but_error: (xsymbol -> 'a -> 'a) -> 'a -> 'a

`feed symbol startp semv endp env` causes the parser to consume the (terminal or nonterminal) symbol `symbol`, accompanied with the semantic value `semv` and with the start and end positions `startp` and `endp`. Thus, the automaton makes a transition, and reaches a new state. The stack grows by one cell. This operation is permitted only if the current state (as determined by `env`) has an outgoing transition labeled with `symbol`. Otherwise, an `Invalid_argument` exception is raised.

      val feed: 'a symbol -> position -> 'a -> position -> 'b env -> 'b env

## 10 Error handling: the traditional way

Menhir’s traditional error handling mechanism is considered deprecated: although it is still supported for the time being, it might be removed in the future. We recommend setting up an error handling mechanism using the new tools offered by Menhir (§11).

#### Error handling

Menhir’s error traditional handling mechanism is inspired by that of yacc and ocamlyacc, but is not identical. A special error token is made available for use within productions. The LR automaton is constructed exactly as if error was a regular terminal symbol. However, error is never produced by the lexical analyzer. Instead, when an error is detected, the current lookahead token is discarded and replaced with the error token, which becomes the current lookahead token. At this point, the parser enters error handling mode.

In error handling mode, automaton states are popped off the automaton’s stack until a state that can act on error is found. This includes both shift and reduce actions. (yacc and ocamlyacc do not trigger reduce actions on error. It is somewhat unclear why this is so.)

When a state that can reduce on error is found, reduction is performed. Since the lookahead token is still error, the automaton remains in error handling mode.

When a state that can shift on error is found, the error token is shifted. At this point, the parser returns to normal mode.

When no state that can act on error is found on the automaton’s stack, the parser stops and raises the exception Error. This exception carries no information. The position of the error can be obtained by reading the lexical analyzer’s environment record.

#### Error recovery

ocamlyacc offers an error recovery mode, which is entered immediately after an error token was successfully shifted. In this mode, tokens are repeatedly taken off the input stream and discarded until an acceptable token is found. This feature is no longer offered by Menhir.

#### Error-related keywords

The following keyword is made available to semantic actions.

When the `$syntaxerror` keyword is evaluated, evaluation of the semantic action is aborted, so that the current reduction is abandoned; the current lookahead token is discarded and replaced with the error token; and error handling mode is entered. Note that there is no mechanism for inserting an error token in front of the current lookahead token, even though this might also be desirable. It is unclear whether this keyword is useful; it might be suppressed in the future.

## 11 Error handling: the new way

Menhir’s incremental API (§9.2) allows taking control when an error is detected. Indeed, as soon as an invalid token is detected, the parser produces a checkpoint of the form `HandlingError _`. At this point, if one decides to let the parser proceed, by just calling `resume`, then Menhir enters its traditional error handling mode (§10). Instead, however, one can decide to take control and perform error handling or error recovery in any way one pleases. One can, for instance, build and display a diagnostic message, based on the automaton’s current stack and/or state. Or, one could modify the input stream, by inserting or deleting tokens, so as to suppress the error, and resume normal parsing. In principle, the possibilities are endless.

An apparently simple-minded approach to error reporting, proposed by Jeffery [10] and further explored by Pottier [20], consists in selecting a diagnostic message (or a template for a diagnostic message) based purely on the current state of the automaton.

In this approach, one determines, ahead of time, which are the “error states” (that is, the states in which an error can be detected), and one prepares, for each error state, a diagnostic message. Because state numbers are fragile (they change when the grammar evolves), an error state is identified not by its number, but by an input sentence that leads to it: more precisely, by an input sentence which causes an error to be detected in this state. Thus, one maintains a set of pairs of an erroneous input sentence and a diagnostic message.

Menhir defines a file format, the .messages file format, for representing this information (§11.1), and offers a set of tools for creating, maintaining, and exploiting .messages files (§11.2). Once one understands these tools, there remains to write a collection of diagnostic messages, a more subtle task than one might think (§11.3), and to glue everything together (§11.4).

In this approach to error handling, as in any other approach, one must understand exactly when (that is, in which states) errors are detected. This in turn requires understanding how the automaton is constructed. Menhir’s construction technique is not Knuth’s canonical LR(1) technique [15], which is usually too expensive to be practical. Instead, Menhir merges states [19] and introduces so-called default reductions. These techniques defer error detection by allowing extra reductions to take place before an error is detected. The impact of these alterations must be taken into account when writing diagnostic messages (§11.3).

In this approach to error handling, the special error token is not used. It should not appear in the grammar. Similarly, the `$syntaxerror` keyword should not be used.

### 11.1 The .messages file format

A .messages file is a text file. Comment lines, which begin with a `#` character, are ignored everywhere. As is evident in the following description, blank lines are significant: they are used as separators between entries and within an entry.

A .messages file is composed of a list of entries. Two entries are separated by one or more blank lines. Each entry consists of one or more input sentences, followed with one or more blank lines, followed with a message. The syntax of an input sentence is described in §8.1. A message is arbitrary text, but cannot contain a blank line. We stress that there cannot be a blank line between two sentences (if there is one, Menhir becomes confused and may complain about some word not being “a known non-terminal symbol”).

>     grammar: TYPE UID
>     grammar: TYPE OCAMLTYPE UID PREC

>     # A (handwritten) comment.

>     Ill-formed declaration.
>     Examples of well-formed declarations:
>       %type <Syntax.expression> expression
>       %type <int> date time

> Figure 16: An entry in a .messages file

>     grammar: TYPE UID
>     ##
>     ## Ends in an error in state: 1.
>     ##
>     ## declaration -> TYPE . OCAMLTYPE separated_nonempty_list(option(COMMA),
>     ##   strict_actual) [ TYPE TOKEN START RIGHT PUBLIC PERCENTPERCENT PARAMETER
>     ##   ON_ERROR_REDUCE NONASSOC LEFT INLINE HEADER EOF COLON ]
>     ##
>     ## The known suffix of the stack is as follows:
>     ## TYPE
>     ##
>     grammar: TYPE OCAMLTYPE UID PREC
>     ##
>     ## Ends in an error in state: 5.
>     ##
>     ## strict_actual -> symbol . loption(delimited(LPAREN,separated_nonempty_list
>     ##   (COMMA,strict_actual),RPAREN)) [ UID TYPE TOKEN START STAR RIGHT QUESTION
>     ##   PUBLIC PLUS PERCENTPERCENT PARAMETER ON_ERROR_REDUCE NONASSOC LID LEFT
>     ##   INLINE HEADER EOF COMMA COLON ]
>     ##
>     ## The known suffix of the stack is as follows:
>     ## symbol
>     ##

>     # A (handwritten) comment.

>     Ill-formed declaration.
>     Examples of well-formed declarations:
>       %type <Syntax.expression> expression
>       %type <int> date time

> Figure 17: An entry in a .messages file, decorated with auto-generated comments

As an example, Figure 16 shows a valid entry, taken from Menhir’s own .messages file. This entry contains two input sentences, which lead to errors in two distinct states. A single message is associated with these two error states.

Several commands, described next (§11.2), produce .messages files where each input sentence is followed with an auto-generated comment, marked with `##`. This special comment indicates in which state the error is detected, and is supposed to help the reader understand what it means to be in this state: What has been read so far? What is expected next?

As an example, the previous entry, decorated with auto-generated comments, is shown in Figure 17. (We have manually wrapped the lines that did not fit in this document.)

An auto-generated comment begins with the number of the error state that is reached via this input sentence.

Then, the auto-generated comment shows the LR(1) items that compose this state, in the same format as in an .automaton file. these items offer a description of the past (that is, what has been read so far) and the future (that is, which terminal symbols are allowed next).

Finally, the auto-generated comment shows what is known about the stack when the automaton is in this state. (This can be deduced from the LR(1) items, but is more readable if shown separately.)

In a canonical LR(1) automaton, the LR(1) items offer an exact description of the past and future. However, in a noncanonical automaton, which is by default what Menhir produces, the situation is more subtle. The lookahead sets can be over-approximated, so the automaton can perform one or more “spurious reductions” before an error is detected. As a result, the LR(1) items in the error state offer a description of the future that may be both incorrect (that is, a terminal symbol that appears in a lookahead set is not necessarily a valid continuation) and incomplete (that is, a terminal symbol that does not appear in any lookahead set may nevertheless be a valid continuation). More details appear further on (§11.3).

In order to attract the user’s attention to this issue, if an input sentence causes one or more spurious reductions, then the auto-generated comment contains a warning about this fact. This mechanism is not completely foolproof, though, as it may be the case that one particular sentence does not cause any spurious reductions (hence, no warning appears), yet leads to an error state that can be reached via other sentences that do involve spurious reductions.

### 11.2 Maintaining .messages files

Ideally, the set of input sentences in a .messages file should be correct (that is, every sentence causes an error on its last token), irredundant (that is, no two sentences lead to the same error state), and complete (that is, every error state is reached by some sentence).

Correctness and irredundancy are checked by the command --compile-errors filename, where filename is the name of a .messages file. This command fails if a sentence does not cause an error at all, or causes an error too early. It also fails if two sentences lead to the same error state. If the file is correct and irredundant, then (as its name suggests) this command compiles the .messages file down to an OCaml function, whose code is printed on the standard output channel. This function, named `message`, has type `int -> string`, and maps a state number to a message. It raises the exception `Not_found` if its argument is not the number of a state for which a message has been defined.

Completeness is checked via the commands --list-errors and --compare-errors. The former produces, from scratch, a complete set of input sentences, that is, a set of input sentences that reaches all error states. The latter compares two sets of sentences (more precisely, the two underlying sets of error states) for inclusion.

The command --list-errors first computes all possible ways of causing an error. From this information, it deduces a list of all error states, that is, all states where an error can be detected. For each of these states, it computes a (minimal) input sentence that causes an error in this state. Finally, it prints these sentences, in the .messages file format, on the standard output channel. Each sentence is followed with an auto-generated comment and with a dummy diagnostic message. The user should be warned that this algorithm may require large amounts of time (typically in the tens of seconds, possibly more) and memory (typically in the gigabytes, possibly more). It requires a 64-bit machine. (On a 32-bit machine, it works, but quickly hits a built-in size limit.) At the verbosity level --log-automaton 2, it displays some progress information and internal statistics on the standard error channel.

The command --compare-errors filename1 --compare-errors filename2 compares the .messages files filename1 and filename2. Each file is read and internally translated to a mapping of states to messages. Menhir then checks that the left-hand mapping is a subset of the right-hand mapping. That is, if a state s is reached by some sentence in filename1, then it should also be reached by some sentence in filename2. Furthermore, if the message associated with s in filename1 is not a dummy message, then the same message should be associated with s in filename2.

To check that the sentences in filename2 cover all error states, it suffices to (1) use --list-errors to produce a complete set of sentences, which one stores in filename1, then (2) use --compare-errors to compare filename1 and filename2.

In the case of a grammar that evolves fairly often, it can take significant human time and effort to update the .messages file and ensure correctness, irredundancy, and completeness. A way of reducing this effort is to abandon completeness. This implies that the auto-generated `message` function can raise `Not_found` and that a generic “syntax error” message must be produced in that case. We prefer to discourage this approach, as it implies that the end user is exposed to a mixture of specific and generic syntax error messages, and there is no guarantee that the specific (hand-written) messages will appear in all situations where there are expected to appear. Instead, we recommend waiting for the grammar to become stable and enforcing completeness.

The command --update-errors filename is used to update the auto-generated comments in the .messages file filename. It is typically used after a change in the grammar (or in the command line options that affect the construction of the automaton). A new .messages file is produced on the standard output channel. It is identical to filename, except the auto-generated comments, identified by `##`, have been removed and re-generated.

The command --echo-errors filename is used to filter out all comments, blank lines, and messages from the .messages file filename. The input sentences, and nothing else, are echoed on the standard output channel. As an example application, one could then translate the sentences to concrete syntax and create a collection of source files that trigger every possible syntax error.

The command --interpret-error is analogous to --interpret. It causes Menhir to act as an interpreter. Menhir reads sentences off the standard input channel, parses them, and displays the outcome. This switch can be usefully combined with --trace. The main difference between --interpret and --interpret-error is that, when the latter command is used, Menhir expects the input sentence to cause an error on its last token, and displays information about the state in which the error is detected, in the form of a .messages file entry. This can be used to quickly find out exactly what error is caused by one particular input sentence.

### 11.3 Writing accurate diagnostic messages

One might think that writing a diagnostic message for each error state is a straightforward (if lengthy) task. In reality, it is not so simple.

#### A state, not a sentence

The first thing to keep in mind is that a diagnostic message is associated with a state s, as opposed to a sentence. An entry in a .messages file contains a sentence w that leads to an error in state s. This sentence is just one way of causing an error in state s; there may exist many other sentences that also cause an error in this state. The diagnostic message should not be specific of the sentence w: it should make sense regardless of how the state s is reached.

As a rule of thumb, when writing a diagnostic message, one should (as much as possible) ignore the example sentence w altogether, and concentrate on the description of the state s, which appears as part of the auto-generated comment.

The LR(1) items that compose the state s offer a description of the past (that is, what has been read so far) and the future (that is, which terminal symbols are allowed next). A diagnostic message should be designed based on this description.

>     %token ID ARROW LPAREN RPAREN COLON SEMICOLON
>     %start<unit> program
>     %%
>     typ0: ID | LPAREN typ1 RPAREN {}
>     typ1: typ0 | typ0 ARROW typ1  {}
>     declaration: ID COLON typ1    {}
>     program:
>     | LPAREN declaration RPAREN
>     | declaration SEMICOLON       {}

> Figure 18: A grammar where one error state is difficult to explain

>     program: ID COLON ID LPAREN
>     ##
>     ## Ends in an error in state: 8.
>     ##
>     ## typ1 -> typ0 . [ SEMICOLON RPAREN ]
>     ## typ1 -> typ0 . ARROW typ1 [ SEMICOLON RPAREN ]
>     ##
>     ## The known suffix of the stack is as follows:
>     ## typ0
>     ##

> Figure 19: A problematic error state in the grammar of Figure 18, due to over-approximation

#### The problem of over-approximated lookahead sets

As pointed out earlier (§11.1), in a noncanonical automaton, the lookahead sets in the LR(1) items can be both over- and under-approximated. One must be aware of this phenomenon, otherwise one runs the risk of writing a diagnostic message that proposes too many or too few continuations.

As an example, let us consider the grammar in Figure 18. According to this grammar, a “program” is either a declaration between parentheses or a declaration followed with a semicolon. A “declaration” is an identifier, followed with a colon, followed with a type. A “type” is an identifier, a type between parentheses, or a function type in the style of OCaml.

The (noncanonical) automaton produced by Menhir for this grammar has 17 states. Using --list-errors, we find that an error can be detected in 10 of these 17 states. By manual inspection of the auto-generated comments, we find that for 9 out of these 10 states, writing an accurate diagnostic message is easy. However, one problematic state remains, namely state 8, shown in Figure 19.

In this state, a (level-0) type has just been read. One valid continuation, which corresponds to the second LR(1) item in Figure 19, is to continue this type: the terminal symbol `ARROW`, followed with a (level-1) type, is a valid continuation. Now, the question is, what other valid continuations are there? By examining the first LR(1) item in Figure 19, it may look as if both `SEMICOLON` and `RPAREN` are valid continuations. However, this cannot be the case. A moment’s thought reveals that either we have seen an opening parenthesis `LPAREN` at the very beginning of the program, in which case we definitely expect a closing parenthesis `RPAREN`; or we have not seen one, in which case we definitely expect a semicolon `SEMICOLON`. It is never the case that both `SEMICOLON` and `RPAREN` are valid continuations!

In fact, the lookahead set in the first LR(1) item in Figure 19 is over-approximated. State 8 in the noncanonical automaton results from merging two states in the canonical automaton.

In such a situation, one cannot write an accurate diagnostic message. Knowing that the automaton is in state 8 does not give us a precise view of the valid continuations. Some valuable information (that is, whether we have seen an opening parenthesis `LPAREN` at the very beginning of the program) is buried in the automaton’s stack.

>     %token ID ARROW LPAREN RPAREN COLON SEMICOLON
>     %start<unit> program
>     %%
>     typ0: ID | LPAREN typ1(RPAREN) RPAREN          {}
>     typ1(phantom): typ0 | typ0 ARROW typ1(phantom) {}
>     declaration(phantom): ID COLON typ1(phantom)   {}
>     program:
>     | LPAREN declaration(RPAREN) RPAREN
>     | declaration(SEMICOLON)  SEMICOLON            {}

> Figure 20: Splitting the problematic state of Figure 19 via selective duplication

>     %token ID ARROW LPAREN RPAREN COLON SEMICOLON
>     %start<unit> program
>     %on_error_reduce typ1
>     %%
>     typ0: ID | LPAREN typ1 RPAREN {}
>     typ1: typ0 | typ0 ARROW typ1  {}
>     declaration: ID COLON typ1    {}
>     program:
>     | LPAREN declaration RPAREN
>     | declaration SEMICOLON       {}

> Figure 21: Avoiding the problematic state of Figure 19 via reductions on error

>     program: ID COLON ID LPAREN
>     ##
>     ## Ends in an error in state: 15.
>     ##
>     ## program -> declaration . SEMICOLON [ # ]
>     ##
>     ## The known suffix of the stack is as follows:
>     ## declaration
>     ##
>     ## WARNING: This example involves spurious reductions.
>     ## This implies that, although the LR(1) items shown above provide an
>     ## accurate view of the past (what has been recognized so far), they
>     ## may provide an INCOMPLETE view of the future (what was expected next).
>     ## In state 8, spurious reduction of production typ1 -> typ0
>     ## In state 11, spurious reduction of production declaration -> ID COLON typ1
>     ##

> Figure 22: A problematic error state in the grammar of Figure 21, due to under-approximation

How can one work around this problem? Let us suggest three options.

#### Blind duplication of states

One option would be to build a canonical automaton by using the --canonical switch. In this example, one would obtain a 27-state automaton, where the problem has disappeared. However, this option is rarely viable, as it duplicates many states without good reason.

#### Selective duplication of states

A second option is to manually cause just enough duplication to remove the problematic over-approximation. In our example, we wish to distinguish two kinds of types and declarations, namely those that must be followed with a closing parenthesis, and those that must be followed with a semicolon. We create such a distinction by parameterizing `typ1` and `declaration` with a phantom parameter. The modified grammar is shown in Figure 20. The phantom parameter does not affect the language that is accepted: for instance, the nonterminal symbols declaration(SEMICOLON) and declaration(RPAREN) generate the same language as declaration in the grammar of Figure 18. Yet, by giving distinct names to these two symbols, we force the construction of an automaton where more states are distinguished. In this example, Menhir produces a 23-state automaton. Using --list-errors, we find that an error can be detected in 11 of these 23 states, and by manual inspection of the auto-generated comments, we find that for each of these 11 states, writing an accurate diagnostic message is easy. In summary, we have selectively duplicated just enough states so as to split the problematic error state into two non-problematic error states.

#### Reductions on error

A third and last option is to introduce an %on_error_reduce declaration (§4.1.8) so as to prevent the detection of an error in the problematic state 8. We see in Figure 19 that, in state 8, the production typ1 → typ0 is ready to be reduced. If we could force this reduction to take place, then the automaton would move to some other state where it would be clear which of `SEMICOLON` and `RPAREN` is expected. We achieve this by marking `typ1` as “reducible on error”. The modified grammar is shown in Figure 21. For this grammar, Menhir produces a 17-state automaton. (This is the exact same automaton as for the grammar of Figure 18, except 2 of the 17 states have received extra reduction actions.) Using --list-errors, we find that an error can be detected in 9 of these 17 states. The problematic state, namely state 8, is no longer an error state! The problem has vanished.

#### The problem of under-approximated lookahead sets

The third option seems by far the simplest of all, and is recommended in many situations. However, it comes with a caveat. There may now exist states whose lookahead sets are under-approximated, in a certain sense. Because of this, there is a danger of writing an incomplete diagnostic message, one that does not list all valid continuations.

To see this, let us look again at the sentence ID COLON ID LPAREN. In the grammar and automaton of Figure 18, this sentence takes us to the problematic state 8, shown in Figure 19. In the grammar and automaton of Figure 21, because more reduction actions are carried out before the error is detected, this sentence takes us to state 15, shown in Figure 22.

When writing a diagnostic message for state 15, one might be tempted to write: “Up to this point, a declaration has been recognized. At this point, a semicolon is expected”. Indeed, by examining the sole LR(1) item in state 15, it looks as if `SEMICOLON` is the only permitted continuation. However, this is not the case. Another valid continuation is `ARROW`: indeed, the sentence ID COLON ID ARROW ID SEMICOLON forms a valid program. In fact, if the first token following ID COLON ID is ARROW, then in state 8 this token is shifted, so the two reductions that take us from state 8 through state 11 to state 15 never take place. This is why, even though ARROW does not appear in state 15 as a valid continuation, it nevertheless is a valid continuation of ID COLON ID. The warning produced by Menhir, shown in Figure 22, is supposed to attract attention to this issue.

Another way to explain this issue is to point out that, by declaring `%on_error_reduce typ1`, we make a choice. When the parser reads a type and finds an invalid token, it decides that this type is finished, even though, in reality, this type could be continued with `ARROW` …. This in turn causes the parser to perform another reduction and consider the current declaration finished, even though, in reality, this declaration could be continued with `ARROW` ….

In summary, when writing a diagnostic message for state 15, one should take into account the fact that this state can be reached via spurious reductions and (therefore) `SEMICOLON` may not be the only permitted continuation. One way of doing this, without explicitly listing all permitted continuations, is to write: “Up to this point, a declaration has been recognized. If this declaration is complete, then at this point, a semicolon is expected”.

### 11.4 A working example

The CompCert verified compiler offers a real-world example of this approach to error handling. The “pre-parser” is where syntax errors are detected: see cparser/pre_parser.mly. A database of erroneous input sentences and (templates for) diagnostic messages is stored in cparser/handcrafted.messages. It is compiled, using --compile-errors, to an OCaml file named cparser/pre_parser_messages.ml. The function `Pre_parser_messages.message`, which maps a state number to (a template for) a diagnostic message, is called from cparser/ErrorReports.ml, where we construct and display a full-fledged diagnostic message.

In CompCert, we allow a template for a diagnostic message to contain the special form `$i`, where `i` is an integer constant, understood as an index into the parser’s stack. The code in cparser/ErrorReports.ml automatically replaces this special form with the fragment of the source text that corresponds to this stack entry. This mechanism is not built into Menhir ; it is implemented in CompCert using Menhir’s incremental API.

## 12 Coq back-end

Menhir is able to generate a parser that whose correctness can be formally verified using the Coq proof assistant [13]. This feature is used to construct the parser of the CompCert verified compiler [17].

Setting the --coq switch on the command line enables the Coq back-end. When this switch is set, Menhir expects an input file whose name ends in .vy and generates a Coq file whose name ends in .v.

Like a .mly file, a .vy file is a grammar specification, with embedded semantic actions. The only difference is that the semantic actions in a .vy file are expressed in Coq instead of OCaml. A .vy file otherwise uses the same syntax as a .mly file. CompCert’s cparser/Parser.vy serves as an example.

Several restrictions are imposed when Menhir is used in --coq mode:

- The error handling mechanism (§10) is absent. The `$syntaxerror` keyword and the error token are not supported.
- Location information is not propagated. The `$start*` and `$end*` keywords (Figure 14) are not supported.
- %parameter (§4.1.2) is not supported.
- %inline (§5.3) is not supported.
- The standard library (§5.4) is not supported, of course, because its semantic actions are expressed in OCaml. If desired, the user can define an analogous library, whose semantic actions are expressed in Coq.
- Because Coq’s type inference algorithm is rather unpredictable, the Coq type of every nonterminal symbol must be provided via a %type or %start declaration (§4.1.5, §4.1.6).
- Unless the proof of completeness has been deactivated using --coq-no-complete, the grammar must not have a conflict (not even a benign one, in the sense of §6.1). That is, the grammar must be LR(1). Conflict resolution via priority and associativity declarations (§4.1.4) is not supported. The reason is that there is no simple formal specification of how conflict resolution should work.

The generated file contains several modules:

- The module `Gram` defines the terminal and non-terminal symbols, the grammar, and the semantic actions.
- The module `Aut` contains the automaton generated by Menhir, together with a certificate that is checked by Coq while establishing the soundness and completeness of the parser.

The type `terminal` of the terminal symbols is an inductive type, with one constructor for each terminal symbol. A terminal symbol named `Foo` in the `.vy` file is named `Foo't` in Coq. A terminal symbol per se does not carry a the semantic value.

We also define the type `token` of tokens, that is, dependent pairs of a terminal symbol and a semantic value of an appropriate type for this symbol. We model the lexer as an object of type `Streams.Stream token`, that is, an infinite stream of tokens.

The type `nonterminal` of the non-terminal symbols is an inductive type, with one constructor for each non-terminal symbol. A non-terminal symbol named `Bar` in the `.vy` file is named `Bar'nt` in Coq.

The proof of termination of an LR(1) parser in the case of invalid input seems far from obvious. We did not find such a proof in the literature. In an application such as CompCert [17], this question is not considered crucial. For this reason, we did not formally establish the termination of the parser. Instead, in order to satisfy Coq’s termination requirements, we use the “fuel” technique: the parser takes an additional parameter `log_fuel` of type `nat` such that 2`log_fuel` is the maximum number of steps the parser is allowed to perform. In practice, one can use a value of e.g., 40 or 50 to make sure the parser will never run out of fuel in a reasonnable time.

Parsing can have three different outcomes, represented by the type `parse_result`. (This definition is implicitly parameterized over the initial state `init`. We omit the details here.)

      Inductive parse_result :=
      | Fail_pr:    parse_result
      | Timeout_pr: parse_result
      | Parsed_pr:
          symbol_semantic_type (NT (start_nt init)) ->
          Stream token ->
          parse_result.

The outcome `Fail_pr` means that parsing has failed because of a syntax error. (If the completeness of the parser with respect to the grammar has been proved, this implies that the input is invalid). The outcome `Timeout_pr` means that the fuel has been exhausted. Of course, this cannot happen if the parser was given an infinite amount of fuel, as suggested above. The outcome `Parsed_pr` means that the parser has succeeded in parsing a prefix of the input stream. It carries the semantic value that has been constructed for this prefix, as well as the remainder of the input stream.

For each entry point `entry` of the grammar, Menhir generates a parsing function `entry`, whose type is `nat -> Stream token -> parse_result`.

Two theorems are provided, named `entry_point_correct` and `entry_point_complete`. The correctness theorem states that, if a word (a prefix of the input stream) is accepted, then this word is valid (with respect to the grammar) and the semantic value that is constructed by the parser is valid as well (with respect to the grammar). The completeness theorem states that if a word (a prefix of the input stream) is valid (with respect to the grammar), then (given sufficient fuel) it is accepted by the parser.

These results imply that the grammar is unambiguous: for every input, there is at most one valid interpretation. This is proved by another generated theorem, named `Parser.unambiguous`.

The parsers produced by Menhir’s Coq back-end must be linked with a Coq library. This library can be installed via the command `opam install coq-menhirlib`.2 The Coq sources of this library can be found in the coq-menhirlib directory of the Menhir repository.

The CompCert verified compiler [17,16] can be used as an example if one wishes to use Menhir to generate a formally verified parser as part of some other project. See in particular the directory cparser.

## 13 Building grammarware on top of Menhir

It is possible to build a variety of grammar-processing tools, also known as “grammarware” [14], on top of Menhir’s front-end. Indeed, Menhir offers a facility for dumping a .cmly file, which contains a (binary-form) representation of the grammar and automaton, as well as a library, MenhirSdk, for (programmatically) reading and exploiting a .cmly file. These facilities are described in §13.1. Furthermore, Menhir allows decorating a grammar with “attributes”, which are ignored by Menhir’s back-ends, yet are written to the .cmly file, thus can be exploited by other tools, via MenhirSdk. Attributes are described in §13.2.

### 13.1 Menhir’s SDK

The command line option --cmly causes Menhir to produce a .cmly file in addition to its normal operation. This file contains a (binary-form) representation of the grammar and automaton. This is the grammar that is obtained after the following steps have been carried out:

- joining multiple .mly files, if necessary;
- eliminating anonymous rules;
- expanding away parameterized nonterminal symbols;
- removing unreachable nonterminal symbols;
- performing OCaml type inference, if the --infer switch is used;
- inlining away nonterminal symbols that are decorated with %inline.

The library MenhirSdk offers an API for reading a .cmly file. The functor MenhirSdk.Cmly_read.Read reads such a file and produces a module whose signature is MenhirSdk.Cmly_api.GRAMMAR. This API is not explained in this document; for details, the reader is expected to follow the above links.

### 13.2 Attributes

Attributes are decorations that can be placed in .mly files. They are ignored by Menhir’s back-ends, but are written to .cmly files, thus can be exploited by other tools, via MenhirSdk.

An attribute consists of a name and a payload. An attribute name is an OCaml identifier, such as cost, or a list of OCaml identifiers, separated with dots, such as my.name. An attribute payload is an OCaml expression of arbitrary type, such as 1 or `"&&"` or `print_int`. Following the syntax of OCaml’s attributes, an attribute’s name and payload are separated with one or more spaces, and are delimited by `[@` and `]`. Thus, `[@cost 1]` and `[@printer print_int]` are examples of attributes.

An attribute can be attached at one of four levels:

1. An attribute can be attached with the grammar. Such an attribute must be preceded with a `%` sign and must appear in the declarations section (§4.1). For example, the following is a valid declaration:

         %[@trace true]

2. An attribute can be attached with a terminal symbol. Such an attribute must follow the declaration of this symbol. For example, the following is a valid declaration of the terminal symbol `INT`:

         %token<int> INT [@cost 0] [@printer print_int]

3. An attribute can be attached with a nonterminal symbol. Such an attribute must appear inside the rule that defines this symbol, immediately after the name of this symbol. For instance, the following is a valid definition of the nonterminal symbol `expr`:

         expr [@default EConst 0]:
           i = INT                  { EConst i }
         | e1 = expr PLUS e2 = expr { EAdd (e1, e2) }

   An attribute can be attached with a parameterized nonterminal symbol:

         option [@default None] (X):
                 { None }
         | x = X { Some x }

   An attribute cannot be attached with a nonterminal symbol that is decorated with the %inline keyword.
4. An attribute can be attached with a producer (§4.2.3), that is, with an occurrence of a terminal or nonterminal symbol in the right-hand side of a production. Such an attribute must appear immediately after the producer. For instance, in the following rule, an attribute is attached with the producer `expr*`:

         exprs:
           LPAREN es = expr* [@list true] RPAREN { es }

As a convenience, it is possible to attach many attributes with many (terminal and nonterminal) symbols in one go, via an %attribute declaration, which must be placed in the declarations section (§4.1). For instance, the following declaration attaches both of the attributes `[@cost 0]` and `[@precious false]` with each of the symbols `INT` and `id`:

      %attribute INT id [@cost 0] [@precious false]

An %attribute declaration can be considered syntactic sugar: it is desugared away in terms of the four forms of attributes presented earlier. (The command line switch --only-preprocess can be used to see how it is desugared.)

If an attribute is attached with a parameterized nonterminal symbol, then, when this symbol is expanded away, the attribute is transmitted to every instance. For instance, in an earlier example, the attribute `[@default None]` was attached with the parameterized symbol `option`. Then, every instance of `option`, such as `option(expr)`, `option(COMMA)`, and so on, inherits this attribute. To attach an attribute with one specific instance only, one can use an %attribute declaration. For instance, the declaration `%attribute option(expr) [@cost 10]` attaches an attribute with the nonterminal symbol `option(expr)`, but not with the symbol `option(COMMA)`.

## 14 Interaction with build systems

This section explains some details of the compilation workflow, including OCaml type inference and its repercussions on dependency analysis (§14.1) and compilation flags (§14.2). This material should be of interest only to authors of build systems who wish to build support for Menhir into their system. Ordinary users should skip this section and use a build system that knows about Menhir, such as dune (preferred) or ocamlbuild.

### 14.1 OCaml type inference and dependency analysis

In an ideal world, the semantic actions in a .mly file should be well-typed according to the OCaml type discipline, and their types should be known to Menhir, which may need this knowledge. (When --inspection is set, Menhir needs to know the OCaml type of every nonterminal symbol.) To address this problem, three approaches exist:

- Ignore the problem and let Menhir run without OCaml type information (§14.1.1).
- Let Menhir obtain OCaml type information by invoking the OCaml compiler (§14.1.2).
- Let Menhir request and receive OCaml type information without invoking the OCaml compiler (§14.1.3).

#### 14.1.1 Running without OCaml type information

The simplest thing to do is to run Menhir without any of the flags described in the following (§14.1.2, §14.1.3). Then, the semantic actions are not type-checked, and their OCaml type is not inferred. (This is analogous to using ocamlyacc.) The drawbacks of this approach are as follows:

- A type error in a semantic action is detected only when the .ml file produced by Menhir is type-checked. The location of the type error, as reported by the OCaml compiler, can be suboptimal.
- Unless a %type declaration for every nonterminal symbol is given, the inspection API cannot be generated, that is, --inspection must be turned off.

#### 14.1.2 Obtaining OCaml type information by calling the OCaml compiler

The second approach is to let Menhir invoke the OCaml compiler so as to type-check the semantic actions and infer their types. This is done by invoking Menhir with the --infer switch, as follows.

--infer. This switch causes the semantic actions to be checked for type consistency before the parser is generated. To do so, Menhir generates a mock .ml file, which contains just the semantic actions, and invokes the OCaml compiler, under the form `ocamlc -i`, so as to type-check this file and infer the types of the semantic actions. Menhir then reads this information and produces real .ml and .mli files.

--ocamlc command. This switch controls how ocamlc is invoked. It allows setting both the name of the executable and the command line options that are passed to it.

One difficulty with the this approach is that the OCaml compiler usually needs to consult a few .cm[iox] files. Indeed, if the .mly file contains a reference to an external OCaml module, say A, then the OCaml compiler typically needs to read one or more files named A.cm[iox].

This implies that these files must have been created first. But how is one supposed to know, exactly, which files should be created first? One must scan the .mly file so as to find out which external modules it depends upon. In other words, a dependency analysis is required. This analysis can be carried out by invoking Menhir with the --depend switch, as follows.

--depend. This switch causes Menhir to generate dependency information for use in conjunction with make. When invoked in this mode, Menhir does not generate a parser. Instead, it examines the grammar specification and prints a list of prerequisites for the targets basename.cm[iox], basename.ml, and basename.mli. This list is intended to be textually included within a Makefile. To produce this list, Menhir generates a mock .ml file, which contains just the semantic actions, invokes ocamldep, and postprocesses its output.

--raw-depend. This switch is analogous to --depend. However, in this case, ocamldep’s output is not postprocessed by Menhir: it is echoed without change. This switch is not suitable for direct use with make ; it is intended for use with omake or ocamlbuild, which perform their own postprocessing.

--ocamldep command. This switch controls how ocamldep is invoked. It allows setting both the name of the executable and the command line options that are passed to it.

#### 14.1.3 Obtaining OCaml type information without calling the OCaml compiler

The third approach is to let Menhir request and receive OCaml type information without allowing Menhir to invoke the OCaml compiler. There is nothing magic about this: to achieve this, Menhir must be invoked twice, and the OCaml compiler must be invoked (by the user, or by the build system) in between. This is done as follows.

--infer-write-query mockfilename. When invoked in this mode, Menhir does not generate a parser. Instead, generates a mock .ml file, named mockfilename, which contains just the semantic actions. Then, it stops.

It is then up to the user (or to the build system) to invoke `ocamlc -i` so as to type-check the mock .ml file and infer its signature. The output of this command should be redirected to some file sigfilename. Then, Menhir can be invoked again, as follows.

--infer-read-reply sigfilename. When invoked in this mode, Menhir assumes that the file sigfilename contains the result of running `ocamlc -i` on the file mockfilename. It reads and parses this file, so as to obtain the OCaml type of every semantic action, then proceeds normally to generate a parser.

This protocol was introduced on 2018/05/23; earlier versions of Menhir do not support it. Its existence can be tested as follows:

--infer-protocol-supported. When invoked with this switch, Menhir immediately terminates with exit code 0. An earlier version of Menhir, which does not support this protocol, would display a help message and terminate with a nonzero exit code.

### 14.2 Compilation flags

The following switches allow querying Menhir so as to find out which compilation flags should be passed to the OCaml compiler and linker.

--suggest-comp-flags. This switch causes Menhir to print a set of suggested compilation flags, and exit. These flags are intended to be passed to the OCaml compilers (ocamlc or ocamlopt) when compiling and linking the parser generated by Menhir. What flags are suggested? In the absence of the --table switch, no flags are suggested. When --table is set, a -I flag is suggested, so as to ensure that MenhirLib is visible to the OCaml compiler.

--suggest-link-flags-byte. This switch causes Menhir to print a set of suggested link flags, and exit. These flags are intended to be passed to ocamlc when producing a bytecode executable. What flags are suggested? In the absence of the --table switch, no flags are suggested. When --table is set, the object file menhirLib.cma is suggested, so as to ensure that MenhirLib is linked in.

--suggest-link-flags-opt. This switch causes Menhir to print a set of suggested link flags, and exit. These flags are intended to be passed to ocamlopt when producing a native code executable. What flags are suggested? In the absence of the --table switch, no flags are suggested. When --table is set, the object file menhirLib.cmxa is suggested, so as to ensure that MenhirLib is linked in.

--suggest-menhirLib. This switch causes Menhir to print (the absolute path of) the directory where MenhirLib was installed.

--suggest-ocamlfind. This switch is deprecated and may be removed in the future. It always prints false.

## 15 Comparison with ocamlyacc

Roughly speaking, Menhir is 90% compatible with ocamlyacc. Legacy ocamlyacc grammar specifications are accepted and compiled by Menhir. The resulting parsers run and produce correct parse trees. However, parsers that explicitly invoke functions in the module Parsing behave slightly incorrectly. For instance, the functions that provide access to positions return a dummy position when invoked by a Menhir parser. Porting a grammar specification from ocamlyacc to Menhir requires replacing all calls to Parsing with new Menhir-specific keywords (§7).

Here is an incomplete list of the differences between ocamlyacc and Menhir. The list is roughly sorted by decreasing order of importance.

- Menhir allows the definition of a nonterminal symbol to be parameterized (§5.2). A formal parameter can be instantiated with a terminal symbol, a nonterminal symbol, or an anonymous rule (§4.2.4). A library of standard parameterized definitions (§5.4), including options, sequences, and lists, is bundled with Menhir. EBNF syntax is supported: the modifiers ?, +, and * are sugar for options, nonempty lists, and arbitrary lists (Figure 2).
- ocamlyacc only accepts LALR(1) grammars. Menhir accepts LR(1) grammars, thus avoiding certain artificial conflicts.
- Menhir’s %inline keyword (§5.3) helps avoid or resolve some LR(1) conflicts without artificial modification of the grammar.
- Menhir explains conflicts (§6) in terms of the grammar, not just in terms of the automaton. Menhir’s explanations are believed to be understandable by mere humans.
- Menhir offers an incremental API (in --table mode only) (§9.2). This means that the state of the parser can be saved at any point (at no cost) and that parsing can later be resumed from a saved state.
- Menhir offers a set of tools for building a (complete, irredundant) set of invalid input sentences, mapping each such sentence to a (hand-written) error message, and maintaining this set as the grammar evolves (§11).
- In --coq mode, Menhir produces a parser whose correctness and completeness with respect to the grammar can be checked by Coq (§12).
- Menhir offers an interpreter (§8) that helps debug grammars interactively.
- Menhir allows grammar specifications to be split over multiple files (§5.1). It also allows several grammars to share a single set of tokens.
- Menhir produces reentrant parsers.
- Menhir is able to produce parsers that are parameterized by OCaml modules.
- ocamlyacc requires semantic values to be referred to via keywords: `$1`, `$2`, and so on. Menhir allows semantic values to be explicitly named.
- Menhir warns about end-of-stream conflicts (§6.4), whereas ocamlyacc does not. Menhir warns about productions that are never reduced, whereas, at least in some cases, ocamlyacc does not.
- Menhir offers an option to typecheck semantic actions before a parser is generated: see --infer.
- ocamlyacc produces tables that are interpreted by a piece of C code, requiring semantic actions to be encapsulated as OCaml closures and invoked by C code. Menhir offers a choice between producing tables and producing code. In either case, no C code is involved.
- Menhir makes OCaml’s standard library module Parsing entirely obsolete. Access to locations is now via keywords (§7). Uses of `raise Parse_error` within semantic actions are deprecated. The function `parse_error` is deprecated. They are replaced with keywords (§10).
- Menhir’s error handling mechanism (§10) is inspired by ocamlyacc’s, but is not guaranteed to be fully compatible. Error recovery, also known as re-synchronization, is not supported by Menhir.
- The way in which severe conflicts (§6) are resolved is not guaranteed to be fully compatible with ocamlyacc.
- Menhir warns about unused %token, %nonassoc, %left, and %right declarations. It also warns about %prec annotations that do not help resolve a conflict.
- Menhir accepts OCaml-style comments.
- Menhir allows %start and %type declarations to be condensed.
- Menhir allows two (or more) productions to share a single semantic action.
- Menhir produces better error messages when a semantic action contains ill-balanced parentheses.
- ocamlyacc ignores semicolons and commas everywhere. Menhir regards semicolons and commas as significant, and allows them, or requires them, in certain well-defined places.
- ocamlyacc allows %type declarations to refer to terminal or non-terminal symbols, whereas Menhir requires them to refer to non-terminal symbols. Types can be assigned to terminal symbols with a %token declaration.

## 16 Questions and Answers

⋄ Is Menhir faster than ocamlyacc? What is the speed difference between menhir and menhir --table? A (not quite scientific) benchmark suggests that the parsers produced by ocamlyacc and menhir --table have comparable speed, whereas those produced by menhir are between 2 and 5 times faster. This benchmark excludes the time spent in the lexer and in the semantic actions.

⋄ How do I write Makefile rules for Menhir? This can a bit tricky. If you must do this, see §14. It is recommended instead to use a build system with built-in support for Menhir, such as dune (preferred) or ocamlbuild.

⋄ How do I use Menhir with ocamlbuild? Pass `-use-menhir` to ocamlbuild. To pass options to Menhir, pass `-menhir "menhir <options>"` to ocamlbuild. To use Menhir’s table-based back-end, pass `-menhir "menhir --table"` to ocamlbuild, and either pass `-package menhirLib` to ocamlbuild or add the tag `package(menhirLib)` in the `_tags` file. To combine multiple .mly files, say `a.mly` and `b.mly`, into a single parser, say `parser.{ml,mli}`, create a file named `parser.mlypack` that contains the module names `A B`. See the directory demos/ocamlbuild for examples. To deal with .messages files (§11), use the rules provided in the file demos/ocamlbuild/myocamlbuild.ml.

⋄ How do I use Menhir with dune? Please use dune version 1.4.0 or newer, as it has appropriate built-in rules for Menhir parsers. In the simplest scenario, where the parser resides in a single source file parser.mly, the dune-project file should contain a “stanza” along the following lines:

    (menhir
      (modules parser)
      (flags --explain --dump)
      (infer true)
    )

Ordinary command line switches, like --explain and --dump, are passed as part of the flags line, as done above. The --infer switch has special status and should not be used directly; instead, write (infer true) or (infer false), as done above. (The default is true.) The --table switch can also be listed as part of the flags line; if you do so, then you must add menhirLib to the list of libraries that your code requires, as in the following example:

    (executable
       (name myexecutable)
       (libraries menhirLib)
    )

The directory demos offers several examples. For more details, see dune’s documentation. To deal with .messages files (§11), please use and adapt the rules found in the file src/stage2/dune.

⋄ Menhir reports more shift/reduce conflicts than ocamlyacc! How come? ocamlyacc sometimes merges two states of the automaton that Menhir considers distinct. This happens when the grammar is not LALR(1). If these two states happen to contain a shift/reduce conflict, then Menhir reports two conflicts, while ocamlyacc only reports one. Of course, the two conflicts are very similar, so fixing one will usually fix the other as well.

⋄ I do not use ocamllex. Is there an API that does not involve lexing buffers? Like ocamlyacc, Menhir produces parsers whose monolithic API (§9.1) is intended for use with ocamllex. However, it is possible to convert them, after the fact, to a simpler, revised API. In the revised API, there are no lexing buffers, and a lexer is just a function from unit to tokens. Converters are provided by the library module MenhirLib.Convert. This can be useful, for instance, for users of ulex, the Unicode lexer generator. Also, please note that Menhir’s incremental API (§9.2) does not mention the type `Lexing.lexbuf`. In this API, the parser expects to be supplied with triples of a token and start/end positions of type `Lexing.position`.

⋄ I need both %inline and non-%inline versions of a non-terminal symbol. Is this possible? Define an %inline version first, then use it to define a non-%inline version, like this:

    %inline ioption(X):  (* nothing *) { None } | x = X { Some x }
             option(X): o = ioption(X) { o }

This can work even in the presence of recursion, as illustrated by the following definition of (reversed, left-recursive, possibly empty) lists:

    %inline irevlist(X):    (* nothing *) { [] } | xs = revlist(X) x = X { x :: xs }
             revlist(X): xs = irevlist(X) { xs }

The definition of `irevlist` is expanded into the definition of `revlist`, so in the end, `revlist` receives its normal, recursive definition. One can then view `irevlist` as a variant of `revlist` that is inlined one level deep.

⋄ Can I ship a generated parser while avoiding a dependency on MenhirLib? Yes. One option is to use the code-based back-end (that is, to not use --table). In this case, the generated parser is self-contained. Another option is to use the table-based back-end (that is, use --table) and include a copy of the files `menhirLib.{ml,mli}` together with the generated parser. The command menhir --suggest-menhirLib will tell you where to find these source files.

⋄ Why is $startpos off towards the left? It seems to include some leading whitespace. Indeed, as of 2015/11/04, the computation of positions has changed so as to match ocamlyacc’s behavior. As a result, $startpos can now appear to be too far off to the left. This is explained in §7. In short, the solution is to use `$symbolstartpos` instead.

⋄ Can I pretty-print a grammar in ASCII, HTML, or LATEX format? Yes. Have a look at obelisk [4].

⋄ Does Menhir support mid-rule actions? Yes. See midrule and its explanation in §5.4.

## 17 Technical background

After experimenting with Knuth’s canonical LR(1) technique [15], we found that it really is not practical, even on today’s computers. For this reason, Menhir implements a slightly modified version of Pager’s algorithm [19], which merges states on the fly if it can be proved that no reduce/reduce conflicts will arise as a consequence of this decision. This is how Menhir avoids the so-called mysterious conflicts created by LALR(1) parser generators [7, section 5.7].

Menhir’s algorithm for explaining conflicts is inspired by DeRemer and Pennello’s [6] and adapted for use with Pager’s construction technique.

By default, Menhir produces code, as opposed to tables. This approach has been explored before [3,9]. Menhir performs some static analysis of the automaton in order to produce more compact code.

When asked to produce tables, Menhir performs compression via first-fit row displacement, as described by Tarjan and Yao [23]. Double displacement is not used. The action table is made sparse by factoring out an error matrix, as suggested by Dencker, Dürre, and Heuft [5].

The type-theoretic tricks that triggered our interest in LR parsers [21] are not implemented in Menhir. In the beginning, we did not implement them because the OCaml compiler did not at the time offer generalized algebraic data types (GADTs). Today, OCaml has GADTs, but, as the saying goes, “if it ain’t broken, don’t fix it”.

The main ideas behind the Coq back-end are described in a paper by Jourdan, Pottier and Leroy [13]. The C11 parser in the CompCert compiler [17] is constructed by Menhir and verified by Coq, following this technique. How to construct a correct C11 parser using Menhir is described by Jourdan and Pottier [12].

The approach to error reports presented in §11 was proposed by Jeffery [10] and further explored by Pottier [20].

## 18 Acknowledgements

Menhir’s interpreter (--interpret) and table-based back-end (--table) were implemented by Guillaume Bau, Raja Boujbel, and François Pottier. The project was generously funded by Jane Street Capital, LLC through the “OCaml Summer Project” initiative.

Frédéric Bour provided motivation and an initial implementation for the incremental API, for the inspection API, for attributes, and for MenhirSdk. Merlin, an emacs mode for OCaml, contains an impressive incremental, syntax-error-tolerant OCaml parser, which is based on Menhir and has been a driving force for Menhir’s APIs.

Jacques-Henri Jourdan designed and implemented the Coq back-end and did the Coq proofs for it.

Gabriel Scherer provided motivation for investigating Jeffery’s technique.

## References

[1]
  Alfred V. Aho, Ravi Sethi, and Jeffrey D. Ullman. Compilers: Principles, Techniques, and Tools. Addison-Wesley, 1986.
[2]
  Andrew Appel. Modern Compiler Implementation in ML. Cambridge University Press, 1998.
[3]
  Achyutram Bhamidipaty and Todd A. Proebsting. Very fast YACC-compatible parsers (for very little effort). Software: Practice and Experience, 28(2):181–190, 1998.
[4]
  Lélio Brun. Obelisk. https://github.com/Lelio-Brun/Obelisk, 2017.
[5]
  Peter Dencker, Karl Dürre, and Johannes Heuft. Optimization of parser tables for portable compilers. ACM Transactions on Programming Languages and Systems, 6(4):546–572, 1984.
[6]
  Frank DeRemer and Thomas Pennello. Efficient computation of LALR(1) look-ahead sets. ACM Transactions on Programming Languages and Systems, 4(4):615–649, 1982.
[7]
  Charles Donnelly and Richard Stallman. Bison, 2015.
[8]
  John E. Hopcroft, Rajeev Motwani, and Jeffrey D. Ullman. Introduction to Automata Theory, Languages, and Computation. Addison-Wesley, 2000.
[9]
  R. Nigel Horspool and Michael Whitney. Even faster LR parsing. Software: Practice and Experience, 20(6):515–535, 1990.
[10]
  Clinton L. Jeffery. Generating LR syntax error messages from examples. ACM Transactions on Programming Languages and Systems, 25(5):631–640, 2003.
[11]
  Steven C. Johnson. Yacc: Yet another compiler compiler. In UNIX Programmer’s Manual, volume 2, pages 353–387. Holt, Rinehart, and Winston, 1979.
[12]
  Jacques-Henri Jourdan and François Pottier. A simple, possibly correct LR parser for C11. ACM Transactions on Programming Languages and Systems, 39(4):14:1–14:36, August 2017.
[13]
  Jacques-Henri Jourdan, François Pottier, and Xavier Leroy. Validating LR(1) parsers. volume 7211, pages 397–416, 2012.
[14]
  Paul Klint, Ralf Lämmel, and Chris Verhoef. Toward an engineering discipline for grammarware. 14(3):331–380, 2005.
[15]
  Donald E. Knuth. On the translation of languages from left to right. Information & Control, 8(6):607–639, 1965.
[16]
  Xavier Leroy. The CompCert C verified compiler. https://github.com/AbsInt/CompCert, 2014.
[17]
  Xavier Leroy. The CompCert C compiler. http://compcert.inria.fr/, 2015.
[18]
  Xavier Leroy, Damien Doligez, Alain Frisch, Jacques Garrigue, Didier Rémy, and Jérôme Vouillon. The OCaml system: documentation and user’s manual, 2016.
[19]
  David Pager. A practical general method for constructing LR(k) parsers. Acta Informatica, 7:249–268, 1977.
[20]
  François Pottier. Reachability and error diagnosis in LR(1) parsers. In Compiler Construction (CC), pages 88–98, 2016.
[21]
  François Pottier and Yann Régis-Gianas. Towards efficient, typed LR parsers. Electronic Notes in Theoretical Computer Science, 148(2):155–180, 2006.
[22]
  David R. Tarditi and Andrew W. Appel. ML-Yacc User’s Manual, 2000.
[23]
  Robert Endre Tarjan and Andrew Chi-Chih Yao. Storing a sparse table. Communications of the ACM, 22(11):606–611, 1979.

1
  The computation of $symbolstartpos is optimized by Menhir under two assumptions about the lexer. First, Menhir assumes that the lexer never produces a token whose start and end positions are equal. Second, Menhir assumes that two positions produced by the lexer are equal if and only if they are physically equal. If the lexer violates either of these assumptions, the computation of $symbolstartpos could produce a result that differs from Parsing.symbol_start_pos().
2
  This assumes that you have installed opam, the OCaml package manager, and that you have run the command opam repo add coq-released https://coq.inria.fr/opam/released.

> This document was translated from LATEX by H E V E A.