You'd do well to review your readings to see if they stipulate somewhere earlier that when they say "random" they mean uniform random (equiprobable).
But even without such a stipulation that I'd say the definition as you've loosely formulated it and are interpreting it seems to imply equiprobability. If we follow your logic strictly, then we have to conclude that your proposed attack on the normally-distributed output of $G$ would imply that $G$ is not in fact a secure PRNG by your definition and application thereof. And in fact you should be able to conclude that your definition and interpretation implies that the output of a secure PRNG must be uniform.
But the definitions used in theoretical cryptography are much more precise than this. For example in Katz & Lindell's textbook (2nd edition), Definition 3.14 (p. 62):
DEFINITION 3.14. Let $\ell$ be a polynomial and let $G$ be a deterministic polynomial-time algorithms such that for any $n$ and any input $ \in \{0,1\}^n$, the result $G(s)$ is a string of length $\ell(n)$. We say that $G$ is a pseudorandom generator if the following conditions hold:
- (Expansion:) For every $n$ it holds that $\ell(n) > n$.
- (Pseudorandomness:) For any PPT algorithm $D$, there is a negligible function $\mathsf{negl}$ such that $$\bigg|\mathrm{Pr}\big[D(G(s)) = 1\big] - \mathrm{Pr}\big[D(r) = 1\big]\bigg| ≤ \mathsf{negl}(n)$$ where thre first probability is taken over uniform choice of $s \in \{0,1\}^n$ and the randomness of $D$, and the second probability is taken over uniform choice of $r \in \{0,1\}^{\ell(n)}$ and the randomness of $D$.
I'll repeat this bit that I boldfaced at the end:
the second probability is taken over uniform choice of $r \in \{0,1\}^{\ell(n)}$
If your reading materials are leaving you with questions like these, maybe have a look at something more rigorous.