The following test grammars differ only in that the first alternative of the rule 'expr' is either specified inline or refers to another rule 'notExpression' with just the same definition. But this grammars produce different trees parsing this: '! a & b'. Why?
I really want the grammar to produce the first result (with NOT associated with identifier, not with AND expression) but still need to have 'expr' to reference 'notExpression' in my real grammar. What do I have to change?
grammar test;
s: expr ';' <EOF>;
expr:
NOT expr
| left=expr AND right=expr
| identifier
;
identifier: LETTER (LETTER)*;
WS : ' '+ ->skip;
NOT: '!';
AND: '&';
LETTER: 'A'..'z';
Tree one
grammar test;
s: expr ';' <EOF>;
expr:
notExpression
| left=expr AND right=expr
| identifier
;
notExpression: NOT expr;
identifier: LETTER (LETTER)*;
WS : ' '+ ->skip;
NOT: '!';
AND: '&';
LETTER: 'A'..'z';
Tree two
I kind of got an answer to the second part of my question, which still do not quite give me a satisfaction because using this approach in real elaborate grammar is going to be ugly. As to the first part (WHY) I still have no idea, so more answers are welcome.
Anyway, to fix precedence in presence of referenced rule the 'notExpression' rule can be modified as follows:
notExpression: NOT (identifier|expr);
Which produces the tree different from both shown in original question, but at least the NOT does acquire higher precedence.
Parse tree
This is part of my parser grammar:
expression:
multiplyingExpression
(
PLUS multiplyingExpression #plus
| MINUS multiplyingExpression #minus
)*;
I want to define plus and minus Alternative Label, apparently it doesn't allow me to do so.
at #plus and #minus it gives me the error:
missing RPAREN at '#' while look for rule element
Anybody knows how to achieve this whithout changing the structure of this rule's definition?
An alt label can only be defined on the outer edge (non-nested) of an alt. Therefore, cannot be done without changing the structure of the rule.
What can be done is to use an ordinary label to effectively achieve the desired result.
expression:
multiplyingExpression
( type+=PLUS multiplyingExpression
| type+=MINUS multiplyingExpression
)*;
The result is that the ExpressionContext will contain List<Token> type; whose successive values will, by inference, identify the alts matched.
i've been working on an antlr4 grammar for Z Notation (ISO UTF version), and the specification calls for a lex phase, and then a "2 phased" parse.
you first lex it into a bunch of NAME (or DECORWORD) tokens, and then you parse the resulting tokens against the operatorTemplate rules in the spec's parser grammar, replace appropriate tokens, and then finally parse your new modified token stream to get the AST.
i have the above working, but i can't figure out how to set the precedence and associativity of the parser rules dynamically, so the parse trees are wrong.
the operator syntax looks like (numbers are precedence):
-generic 5 rightassoc (_ → _)
-function 65 rightassoc (_ ◁ _)
i don't see any api to set the associativity on a rule, so i tried with semantic predicates, something like:
expression:
: {ZSupport.isLeftAssociative()}? expression I expression
| <assoc=right> expression i expression
;
or
expression:
: {ZSupport.isLeftAssociative()}? expression i expression
| <assoc=right> {ZSupport.isRightAssociative()}? expression I expression
;
but then i get "The following sets of rules are mutually left-recursive [expression]"
can this be done?
I was able to accomplish this by moving the semantic predicate:
expression:
: expression {ZSupport.isLeftAssociative()}? I expression
| <assoc=right> expression I expression
;
I was under the impression that this wasn't going to work based on this discussion:
https://stackoverflow.com/a/23677069/7711235
...but it does seem to work correctly in all my test cases...
The sumScalarOperator gives me this error, it seems that antlr see it like a possible infinite recursion loop. How can i avoid it?
sumScalarOperator: function SUM_TOKEN function;
function :
| INTEGER_TOKEN
| NUMERIC_TOKEN
| sumScalarOperator
| ID;
ID : [A-Za-z_-] [a-zA-Z0-9_-]*;
INTEGER_TOKEN: [0-9]+;
NUMERIC_TOKEN: [0-9]+'.'[0-9]+ ;
ANTLR4 can't cope with mutually left-recursive rules, but it can rewrite single left-recursive rules automatically to eliminate the left-recursion, so you can just feed it with something like:
function : function SUM_TOKEN function # sumScalarOperator
| INTEGER_TOKEN # value
| NUMERIC_TOKEN # value
| ID # value
;
Replace the value label with anything you need.
I am doing a parser in bison/flex.
This is part of my code:
I want to implement the assignment production, so the identifier can be both boolean_expr or expr, its type will be checked by a symbol table.
So it allows something like:
int a = 1;
boolean b = true;
if(b) ...
However, it is reduce/reduce if I include identifier in both term and boolean_expr, any solution to solve this problem?
Essentially, what you are trying to do is to inject semantic rules (type information) into your syntax. That's possible, but it is not easy. More importantly, it's rarely a good idea. It's almost always best if syntax and semantics are well delineated.
All the same, as presented your grammar is unambiguous and LALR(1). However, the latter feature is fragile, and you will have difficulty maintaining it as you complete the grammar.
For example, you don't include your assignment syntax in your question, but it would
assignment: identifier '=' expr
| identifier '=' boolean_expr
;
Unlike the rest of the part of the grammar shown, that production is ambiguous, because:
x = y
without knowing anything about y, y could be reduced to either term or boolean_expr.
A possibly more interesting example is the addition of parentheses to the grammar. The obvious way of doing that would be to add two productions:
term: '(' expr ')'
boolean_expr: '(' boolean_expr ')'
The resulting grammar is not ambiguous, but it is no longer LALR(1). Consider the two following declarations:
boolean x = (y) < 7
boolean x = (y)
In the first one, y must be an int so that (y) can be reduced to a term; in the second one y must be boolean so that (y) can be reduced to a boolean_expr. There is no ambiguity; once the < is seen (or not), it is entirely clear which reduction to choose. But < is not the lookahead token, and in fact it could be arbitrarily distant from y:
boolean x = ((((((((((((((((((((((y...
So the resulting unambiguous grammar is not LALR(k) for any k.
One way you could solve the problem would be to inject the type information at the lexical level, by giving the scanner access to the symbol table. Then the scanner could look a scanned identifier token in the symbol table and use the information in the symbol table to decide between one of three token types (or more, if you have more datatypes): undefined_variable, integer_variable, and boolean_variable. Then you would have, for example:
declaration: "int" undefined_variable '=' expr
| "boolean" undefined_variable '=' boolean_expr
;
term: integer_variable
| ...
;
boolean_expr: boolean_variable
| ...
;
That will work but it should be obvious that this is not scalable: every time you add a type, you'll have to extend both the grammar and the lexical description, because the now the semantics is not only mixed up with the syntax, it has even gotten intermingled with the lexical analysis. Once you let semantics out of its box, it tends to contaminate everything.
There are languages for which this really is the most convenient solution: C parsing, for example, is much easier if typedef names and identifier names are distinguished so that you can tell whether (t)*x is a cast or a multiplication. (But it doesn't work so easily for C++, which has much more complicated name lookup rules, and also much more need for semantic analysis in order to find the correct parse.)
But, honestly, I'd suggest that you do not use C -- and much less C++ -- as a model of how to design a language. Languages which are hard for compilers to parse are also hard for human beings to parse. The "most vexing parse" continues to be a regular source of pain for C++ newcomers, and even sometimes trips up relatively experienced programmers:
class X {
public:
X(int n = 0) : data_is_available_(n) {}
operator bool() const { return data_is_available_; }
// ...
private:
bool data_is_available_;
// ...
};
X my_x_object();
// ...
if (!x) {
// This code is unreachable. Can you see why?
}
In short, you're best off with a language which can be parsed into an AST without any semantic information at all. Once the parser has produced the AST, you can do semantic analyses in separate passes, one of which will check type constraints. That's far and away the cleanest solution. Without explicit typing, the grammar is slightly simplified, because an expr now can be any expr:
expr: conjunction | expr "or" conjunction ;
conjunction: comparison | conjunction "and" comparison ;
comparison: product | product '<' product ;
product: factor | product '*' factor ;
factor: term | factor '+' term ;
term: identifier
| constant
| '(' expr ')'
;
Each action in the above would simply create a new AST node and set $$ to the new node. At the end of the parse, the AST is walked to verify that all exprs have the correct type.
If that seems like overkill for your project, you can do the semantic checks in the reduction actions, effectively intermingling the AST walk with the parse. That might seem convenient for immediate evaluation, but it also requires including explicit type information in the parser's semantic type, which adds unnecessary overhead (and, as mentioned, the inelegance of letting semantics interfere with the parser.) In that case, every action would look something like this:
expr : expr '+' expr { CheckArithmeticCompatibility($1, $3);
$$ = NewArithmeticNode('+', $1, $3);
}